Apparatus and method for controlling access to a memory

ABSTRACT

The present invention provides a data processing apparatus and method for controlling access to a memory in the data processing apparatus. The apparatus comprises a processor operable in a plurality of modes and a plurality of domains, said plurality of domains comprising a secure domain and a non-secure domain, said plurality of modes including at least one non-secure mode being a mode in the non-secure domain and at least one secure mode being a mode in the secure domain. The processor is operable such that when executing a program in a secure mode the program has access to secure data which is not accessible when the processor is operating in a non-secure mode. A memory is operable to store data required by the processor and comprises secure memory for storing secure data and non-secure memory for storing non-secure data, the processor being operable to issue a memory access request when access to an item of data in the memory is required. At least one memory management unit is provided which is operable, upon receipt of the memory access request from the processor, to perform conversion of a virtual address specified by the memory access request to a physical address. A first set of tables is provided, each table in the first set containing a number of first descriptors, each first descriptor containing at least a virtual address portion and a corresponding intermediate address portion, and a second set of tables is also provided, with each table in the second set containing a number of second descriptors, each second descriptor containing at least an intermediate address portion and a corresponding physical address portion. The second set of tables are managed by the processor when operating in a privileged mode which is not a non-secure mode, and hence remains secure. The at least one memory management unit is then operable to cause predetermined tables in the first and second set to be referenced to enable the conversion of the virtual address specified by the memory access request to a physical address.

BACKGROUND OF THE INVENTION

[0001] 1. Field of the Invention

[0002] The present invention relates to an apparatus and method forcontrolling access to a memory.

[0003] 2. Description of the Prior Art

[0004] A data processing apparatus will typically include a processorfor running applications loaded onto the data processing apparatus. Theprocessor will operate under the control of an operating system. Thedata required to run any particular application will typically be storedwithin a memory of the data processing apparatus. It will be appreciatedthat the data may consist of the instructions contained within theapplication and/or the actual data values used during the execution ofthose instructions on the processor.

[0005] There arise many instances where the data used by at least one ofthe applications is sensitive data that should not be accessible byother applications that can be run on the processor. An example would bewhere the data processing apparatus is a smart card, and one of theapplications is a security application which uses sensitive data, suchas for example secure keys, to perform validation, authentication,decryption and the like. It is clearly important in such situations toensure that such sensitive data is kept secure so that it cannot beaccessed by other applications that may be loaded onto the dataprocessing apparatus, for example hacking applications that have beenloaded onto the data processing apparatus with the purpose of seeking toaccess that secure data.

[0006] In known systems, it has typically been the job of the operatingsystem developer to ensure that the operating system provides sufficientsecurity to ensure that the secure data of one application cannot beaccessed by other applications running under the control of theoperating system. However, as systems become more complex, the generaltrend is for operating systems to become larger and more complex, and insuch situations it becomes increasingly difficult to ensure sufficientsecurity within the operating system itself.

[0007] Examples of systems seeking to provide secure storage ofsensitive data and to provide protection against malicious program codeare those described in U.S. patent application Ser. No. 2002/0007456 A1and U.S. Pat. No. 6,282,657B and U.S. Pat. No. 6,292,874B.

[0008] Accordingly, it is desirable to provide an improved technique forseeking to retain the security of such secure data contained within thememory of the data processing apparatus.

SUMMARY OF THE INVENTION

[0009] Viewed from a first aspect, the present invention provides a dataprocessing apparatus, comprising: a processor operable in a plurality ofmodes and a plurality of domains, said plurality of domains comprising asecure domain and a non-secure domain, said plurality of modes includingat least one non-secure mode being a mode in the non-secure domain, andat least one secure mode being a mode in the secure domain, saidprocessor being operable such that when executing a program in a securemode said program has access to secure data which is not accessible whensaid processor is operating in a non-secure mode; a memory operable tostore data required by the processor and comprising secure memory forstoring secure data and non-secure memory for storing non-secure data,the processor being operable to issue a memory access request whenaccess to an item of data in the memory is required; at least one memorymanagement unit operable, upon receipt of the memory access request fromthe processor, to perform conversion of a virtual address specified bythe memory access request to a physical address; a first set of tables,each table in the first set containing a number of first descriptors,each first descriptor containing at least a virtual address portion anda corresponding intermediate address portion; a second set of tables,each table in the second set containing a number of second descriptors,each second descriptor containing at least an intermediate addressportion and a corresponding physical address portion, the second set oftables being managed by the processor when operating in a privilegedmode which is not a non-secure mode; the at least one memory managementunit being operable to cause predetermined tables in said first andsecond set to be referenced to enable the conversion of the virtualaddress specified by the memory access request to a physical address.

[0010] In accordance with the present invention, the processor isarranged to be operable in a plurality of modes and a plurality ofdomains, including at least one non-secure mode being a mode in thenon-secure domain, and at least one secure mode being a mode in thesecure domain. Further, a memory is provided for storing data requiredby the processor, the memory consisting of secure memory for storingsecure data and non-secure memory for storing non-secure data. It willbe appreciated that there will not necessarily be any physicaldistinction between the secure memory regions and non-secure memoryregions of the memory at the time of fabrication, and that these regionsmay instead be defined by a secure operating system of the dataprocessing apparatus when operating in the secure domain. Hence, anyphysical part of the memory device may be allocated as secure memory,and any physical part may be allocated as non-secure memory.

[0011] When the processor wishes to access an item of data in thememory, it will issue a memory access request, and at least one memorymanagement unit (MMU) is then used to perform conversion of a virtualaddress specified by the memory access request to a physical address.

[0012] Since the view of memory when in the secure domain differs fromthe view of memory when in the non-secure domain, different conversionsfrom virtual address to physical address will typically be requireddependent on the domain in which the processor is operating. This issuecan be addressed by providing a table for the secure domain and a tablefor the non-secure domain, with each table containing a number ofdescriptors, and each descriptor associating a virtual address portionwith a corresponding physical address portion.

[0013] However, this means that there are certain physical addresseswhich are not accessible to particular domain(s), and these gaps wouldbe apparent to the operating system used in those domain(s). Whilst theoperating system used in the secure domain will have knowledge of thenon-secure domain, and hence will not be concerned by this, theoperating system in the non-secure domain should ideally not need tohave any knowledge of the presence of the secure domain, but insteadshould operate as though the secure domain were not there.

[0014] In order to overcome this problem, the data processing apparatusof the present invention includes a first set of tables, each table inthe first set containing a number of first descriptors, with each firstdescriptor containing at least a virtual address portion and acorresponding intermediate address portion, and a second set of tables,with each table in the second set containing a number of seconddescriptors, each second descriptor containing at least an intermediateaddress portion and a corresponding physical address portion. The secondset of tables are managed by the processor when operating in aprivileged mode which is not a non-secure mode, and hence can beconsidered as secure. By this approach, the operating system running ineither domain sees as its physical address space an intermediate addressspace, and the intermediate address space can be defined for each domainvia the second set of tables. This hence allows complete control of theenvironment of the non-secure operating system used in the non-securedomain, since the secure memory regions can be completely hidden fromthe non-secure operating system's view of its “physical address” space,since it sees as its physical address space the intermediate addressspace, which can be arranged to have a contiguous sequence ofintermediate addresses.

[0015] A further problem exhibited if only a single set of tables isused, rather than the first and second sets envisaged by the presentinvention, is that, given the different views of memory between thenon-secure domain and the secure domain, the physical address space forthe memory accessible in one domain can start at zero, but this will nottypically be the case for the other domain. In that domain, the factthat the physical address space does not start at zero will be evidentto the operating system being used in that domain, and this would hencerequire that operating system to have a certain knowledge about theother domain. However, use of the present invention gives the operatingsystem in either domain the ability to see its own memory space startingfrom zero, which is what an operating system would typically expect.

[0016] Hence, in accordance with the present invention, when the atleast one memory management unit needs to determine a physical addressfor a virtual address specified by a memory access request, it willcause predetermined tables in the first and second set to be referencedto enable the conversion of the virtual address to a physical address.The use of such a two-level memory protection technique enables thecomplexity introduced by the segmentation of the physical address spaceinto secure memory and non-secure memory to be shielded from theoperating systems running in each domain, and thereby enables thephysical address space to be segmented as desired.

[0017] This technique also facilitates easier manipulation of memoryregions within the physical address space, since it is now possible toswap a region of memory between non-secure memory and secure memorymerely by manipulation of the relevant second descriptor(s) in thesecond set of tables.

[0018] In one embodiment, the privileged mode in which the processormust be operating in order to manage the second set of tables is amonitor mode in which the processor is operable to manage the switchingbetween the secure domain and the non-secure domain. The monitor mode isa highly secure mode of operation, and accordingly is a suitable modefrom which to manage the second set of tables.

[0019] In an alternative embodiment, the privileged mode is a privilegedsecure mode. More particularly, in one embodiment, in said at least onenon-secure mode the processor is operable under the control of anon-secure operating system and in said at least one secure mode theprocessor is operable under the control of a secure operating system,the secure operating system being operable to manage the second set oftables when the processor is operating in the privileged secure mode.

[0020] Each of the first and second set of tables may include one ormore tables, but preferably each set of tables includes at least asecure table and a non-secure table. With regards to the second set oftables, all of the tables within that set, whether a secure table or anon-secure table, are placed within secure memory. Whether anyparticular part of memory is secure or non-secure is defined by thelayout of the non-secure tables in the second set, secure memory beingthat memory which is not accessible from the non-secure tables in thesecond set.

[0021] In one embodiment, when switching between the secure domain andthe non-secure domain, the processor is operable in the monitor mode toselect the predetermined tables in said first and second sets, dependenton whether the domain being switched to is the secure domain or thenon-secure domain. Hence, if the domain being switched to is the securedomain, then a secure table within the first set and a secure tablewithin the second set would typically be selected by the processor whenoperating in the monitor mode. Similarly, if the domain being switchedto is the non-secure domain, then non-secure tables will be chosen fromthe first and second set.

[0022] As an alternative to selecting the tables at the time ofswitching between the secure domain and the non-secure domain, thepredetermined tables within the first and second sets may be selectedwhen the tables are to be referenced dependent on whether the processoris operating in a secure mode or a non-secure mode at the time thememory access request is issued.

[0023] In one embodiment, the at least one memory management unitcomprises a first memory management unit and a second memory managementunit. In one particular embodiment, the tables in the first set areassociated with the first memory management unit and the tables in thesecond set are associated with the second memory management unit.

[0024] In such embodiments, if the first memory management unit needs toaccess a first descriptor within a predetermined table of said firstset, it issues a table lookup request specifying an intermediate addressfor that first descriptor, the second memory management unit beingoperable to receive the table lookup request and determine the physicaladdress corresponding to that intermediate address. Hence, theintermediate address generated by the first memory management unit ispassed to the second memory management unit for resolving into aphysical address.

[0025] Preferably, the second memory management unit is then operable tocause the first descriptor at that physical address to be retrieved andreturned to the first memory management unit.

[0026] In one embodiment, the first memory management unit comprises afirst internal storage unit for storing first descriptors retrieved fromthe predetermined table of the first set, and used by the first memorymanagement unit to derive access control information used to perform theconversion of the virtual address into a corresponding intermediateaddress.

[0027] It will be appreciated that the first internal storage unit cantake a variety of forms. However, in one embodiment, the first internalstorage unit is a first translation lookaside buffer (TLB) operable tostore the first descriptors retrieved from the predetermined table ofthe first set.

[0028] It will be appreciated that in one embodiment, the first internalstorage may solely comprise the above first TLB. However, in oneparticular embodiment, the first TLB is a first main TLB for storing thefirst descriptors retrieved by the first memory management unit from thepredetermined table of the first set, and the internal storage furthercomprises a micro-TLB for storing the access control information derivedfrom the first descriptors, the access control information comprisingconversions between a number of virtual address portions andcorresponding intermediate address portions, and the access controlinformation being transferred from the first main TLB to the micro-TLBprior to use of that access control information by the first memorymanagement unit.

[0029] In an alternative embodiment where two memory management unitsare still employed, but the first memory management unit is arranged toperform virtual to physical address translation, the first memorymanagement unit comprises a first internal storage unit for storing newdescriptors derived from corresponding first and second descriptorsretrieved from the predetermined tables of the first and second sets,and used by the first memory management unit to derive access controlinformation used to perform the conversion of the virtual address into acorresponding physical address.

[0030] In one such embodiment, the first internal storage unit is afirst translation lookaside buffer (TLB) operable to store the newdescriptors derived from corresponding first and second descriptors.

[0031] In one particular embodiment, the first TLB is a first main TLBfor storing the new descriptors derived from corresponding first andsecond descriptors, and the internal storage unit further comprises amicro-TLB for storing the access control information, the access controlinformation comprising conversions between a number of virtual addressportions and corresponding physical address portions, and the accesscontrol information being transferred from the first main TLB to themicro-TLB prior to use of that access control information by the firstmemory management unit.

[0032] It will be appreciated that irrespective of how the first memorymanagement unit is organised, the second memory management unit can takea variety of forms. However, in one embodiment, the second memorymanagement unit comprises a second internal storage unit for storingsecond descriptors retrieved from the predetermined table of the secondset, and used by the second memory management unit to derive accesscontrol information used to perform the conversion of the intermediateaddress into a corresponding physical address.

[0033] In one embodiment, the second internal storage unit is a secondtranslation lookaside buffer (TLB) operable to store the seconddescriptors retrieved from the predetermined table of the second set.

[0034] In one particular embodiment, the second TLB may be a second mainTLB for storing the second descriptors retrieved by the second memorymanagement unit from the predetermined table of the second set, and theinternal storage further comprises a micro-TLB for storing the accesscontrol information derived from the second descriptors, the accesscontrol information comprising conversions between a number ofintermediate address portions and corresponding physical addressportions, and the access control information being transferred from thesecond main TLB to the micro-TLB prior to use of that access controlinformation by the second memory management unit.

[0035] In preferred embodiments, the first and second sets of tableseach comprise at least a secure table and a non-secure table, the firstand second TLBs comprising a flag associated with each descriptor storedtherein to identify whether that descriptor is derived from saidnon-secure table or said secure table.

[0036] By the use of such a flag, the main TLB within each MMU does notneed to be flushed every time the processor switches between a securemode and a non-secure mode, or vice versa, and accordingly performanceis not adversely affected as a result of the provision of the non-securetable and the secure table. This is advantageous whether or not theinternal storage further includes a micro-TLB.

[0037] However, in embodiments where a micro-TLB is also provided, themicro-TLB of both the first and second memory management units ispreferably flushed whenever the mode of operation of the processorchanges between a secure mode and a non-secure mode, in the secure modeaccess control information only being transferred to the micro-TLB froma descriptor in the associated first or second main TLB that saidassociated flag indicates is from the secure table, and in thenon-secure mode access control information only being transferred to themicro-TLB from a descriptor in the associated first or second main TLBthat said associated flag indicates is from the non-secure table.Clearly if only one of the memory management units contains a micro-TLB,then only that memory management unit will be subjected to such aflushing operation when the mode of operation of the processor changesbetween a secure mode and a non-secure mode.

[0038] Since the memory management units reference their respectivemicro-TLBs when performing the conversion of the virtual addressspecified by the memory access request into a physical address, thisapproach ensures that only information derived from descriptors in thenon-secure tables can be used to perform such conversion when theprocessor is operating in a non-secure mode.

[0039] In certain embodiments a single non-secure table may be providedwithin the first set of tables which is used for all processes run bythe processor in a non-secure mode, and a single secure table may beprovided within the first set of tables for all processes run by theprocessor in a secure mode. However, in alternative embodiments, therewill be plurality of non-secure tables within the first set, eachnon-secure table containing descriptors pertaining to an associatedprocess executable on the processor, and similarly there will be aplurality of secure tables, with each secure table containingdescriptors pertaining to an associated process executable on theprocessor. In such embodiments, the first main TLB within the firstmemory management unit preferably comprises a process ID fieldassociated with each descriptor stored within the first main TLB toidentify the associated process to which that descriptor pertains. Byincluding multiple non-secure tables and multiple secure tables, thissignificantly improves the flexibility of mappings between virtual andphysical addresses, since different mappings to the intermediate addressspace can be performed for different processes, not just for non-securemodes or secure modes.

[0040] In one embodiment, the second set of tables includes a singlesecure table and a single non-secure table, although more than onesecure table and more than one non-secure table in the second set oftables can be provided if desired.

[0041] As an alternative to providing two separate memory managementunits, the at least one memory management unit may comprise a singlememory management unit, in which case the processor may be operable toexecute table merging code to reference the predetermined tables of thefirst and second sets in order to produce from a first descriptor and anassociated second descriptor a new descriptor associating a virtualaddress portion with a corresponding physical address portion.

[0042] In one embodiment, the table merging code is operable to retrievethe first descriptor after referencing the predetermined table in thesecond set to obtain the physical address of the first descriptor.

[0043] In one embodiment, the table merging code is operable to use thefirst descriptor to determine the intermediate address corresponding tothe virtual address specified by the memory access request, and to thenreference the predetermined table in the second set to obtain the seconddescriptor providing a physical address for that intermediate address,whereafter the table merging code is operable to merge the first andsecond descriptors to produce the new descriptor.

[0044] It will be appreciated that the single memory management unit ofsuch embodiments may take a variety of forms. However, in oneembodiment, the single memory management unit comprises an internalstorage unit for storing the new descriptor produced by the tablemerging code, and used by the single memory management unit to deriveaccess control information used to perform the conversion of the virtualaddress into a corresponding physical address.

[0045] Preferably, the processor is operable to execute the tablemerging code when the access control information required to determinethe physical address for the memory access request is not found in theinternal storage unit.

[0046] The internal storage unit may take a variety of forms. However,in one embodiment the internal storage unit is a translation lookasidebuffer (TLB) operable to store the new descriptors produced by the tablemerging code.

[0047] More particularly, in one embodiment the TLB is a main TLB forstoring the new descriptors obtained by the single memory managementunit from the table merging code, and the internal storage furthercomprises a micro-TLB for storing the access control information derivedfrom the new descriptors, the access control information comprisingconversions between a number of virtual address portions andcorresponding physical address portions, and the access controlinformation being transferred from the main TLB to the micro-TLB priorto use of that access control information by the single memorymanagement unit.

[0048] In one embodiment, the first and second sets of tables eachcomprise at least a secure table and a non-secure table, the TLBcomprising a flag associated with each new descriptor stored therein toidentify whether that new descriptor is derived from said non-securetables or said secure tables.

[0049] In such embodiments, the micro-TLB of the single memorymanagement unit is preferably flushed whenever the mode of operation ofthe processor changes between a secure mode and a non-secure mode, inthe secure mode access control information only being transferred to themicro-TLB from a new descriptor in the main TLB that said associatedflag indicates is derived from secure tables, and in the non-secure modeaccess control information only being transferred to the micro-TLB froma new descriptor in the main TLB that said associated flag indicates isderived from non-secure tables.

[0050] It will be appreciated that the table merging code may beexecuted by the processor in any predetermined mode. However, in oneembodiment, the table merging code is executed by the processor whenoperating in the monitor mode. The monitor mode of preferred embodimentsemploys a flat mapping, i.e. memory access requests issued by theprocessor when in the monitor mode directly specify physical addresses,and this property of the monitor mode has been found to enable moreready implementation of the table merging code.

[0051] As will be appreciated by those skilled in the art, the tables ofdescriptors of the first and second set may take a variety of forms, butin preferred embodiments such tables are page tables defining for eachof a number of pages of memory a corresponding descriptor within whichis described access control information relevant to that page of memory.Hence, a first descriptor may for example define a virtual addressportion and a corresponding intermediate address portion for the page,access permission information, such as whether the page is accessible insupervisor mode, user mode, etc, and region attributes such as whetherthe data contained within that page is cacheable or bufferable, etc.Similarly a second descriptor may for example define an intermediateaddress portion and a corresponding physical address portion for thepage, and access permission information, such as whether that page isaccessible in supervisor mode, user mode, etc, and region attributessuch as whether the data contained within that page is cacheable, orbufferable, etc. It will be appreciated that suitable conflictresolution processes could be provided in the event that the accesspermission information or region attribute information obtained from afirst descriptor conflicts with the corresponding information in thesecond descriptor. Alternatively, one of the two descriptors could bearranged not to provide such access permission and region attributesinformation.

[0052] It will be appreciated that the first and second sets of tablesmay be stored at any appropriate location within the data processingapparatus, but preferably are stored within the memory of the dataprocessing apparatus.

[0053] Viewed from a second aspect, the present invention provides amethod of controlling access to a memory in a data processing apparatus,the data processing apparatus comprising a processor operable in aplurality of modes and a plurality of domains, said plurality of domainscomprising a secure domain and a non-secure domain, said plurality ofmodes including at least one non-secure mode being a mode in thenon-secure domain, and at least one secure mode being a mode in thesecure domain, said processor being operable such that when executing aprogram in a secure mode said program has access to secure data which isnot accessible when said processor is operating in a non-secure mode,the memory being operable to store data required by the processor andcomprising secure memory for storing secure data and non-secure memoryfor storing non-secure data, the method comprising the steps of:providing a first set of tables, each table in the first set containinga number of first descriptors, each first descriptor containing at leasta virtual address portion and a corresponding intermediate addressportion; providing a second set of tables, each table in the second setcontaining a number of second descriptors, each second descriptorcontaining at least an intermediate address portion and a correspondingphysical address portion, the second set of tables being managed by theprocessor when operating in a privileged mode which is not a non-securemode; issuing from the processor a memory access request when access toan item of data in the memory is required; and performing conversion ofa virtual address specified by the memory access request to a physicaladdress with reference to predetermined tables in said first and secondset.

[0054] Viewed from a third aspect, the present invention provides acomputer program providing table merging code and operable to configurea processor of a data processing apparatus to perform the step of:referencing the predetermined tables of the first and second sets inorder to produce from a first descriptor and an associated seconddescriptor a new descriptor associating a virtual address portion with acorresponding physical address portion. A fourth aspect of the presentinvention provides a computer program product carrying such a computerprogram.

BRIEF DESCRIPTION OF THE DRAWINGS

[0055] The present invention will be described further, by way ofexample only, with reference to preferred embodiments thereof asillustrated in the accompanying drawings, in which:

[0056]FIG. 1 is a block diagram schematically illustrating a dataprocessing apparatus in accordance with preferred embodiments of thepresent invention;

[0057]FIG. 2 schematically illustrates different programs operating in anon-secure domain and a secure domain;

[0058]FIG. 3 schematically illustrates a matrix of processing modesassociated with different security domains;

[0059]FIGS. 4 and 5 schematically illustrate different relationshipsbetween processing modes and security domains;

[0060]FIG. 6 illustrates one programmer's model of a register bank of aprocessor depending upon the processing mode;

[0061]FIG. 7 illustrates an example of providing separate register banksfor a secure domain and a non-secure domain;

[0062]FIG. 8 schematically illustrates a plurality of processing modeswith switches between security domains being made via a separate monitormode;

[0063]FIG. 9 schematically illustrates a scenario for security domainswitching using a mode switching software interrupt instruction;

[0064]FIG. 10 schematically illustrates one example of how non-secureinterrupt requests and secure interrupt requests may be processed by thesystem;

[0065]FIGS. 11A and 11B schematically illustrate an example ofnon-secure interrupt request processing and an example of secureinterrupt request processing in accordance with FIG. 10;

[0066]FIG. 12 illustrates an alternative scheme for the handling ofnon-secure interrupt request signals and secure interrupt requestsignals compared to that illustrated in FIG. 10;

[0067]FIGS. 13A and 13B illustrate example scenarios for dealing with anon-secure interrupt request and a secure interrupt request inaccordance with the scheme illustrated in FIG. 12;

[0068]FIG. 14 is an example of a vector interrupt table;

[0069]FIG. 15 schematically illustrates multiple vector interrupt tablesassociated with different security domains;

[0070]FIG. 16 schematically illustrates an exception control register;

[0071]FIG. 17 is a flow diagram illustrating how an instructionattempting to change a processing status register in a manner thatalters the security domain setting can generate a separate mode changeexception which in turn triggers entry into the monitor mode and runningof the monitor program;

[0072]FIG. 18 schematically shows a thread of control of a processoroperating in a plurality of modes, wherein a task in monitor mode isinterrupted;

[0073]FIG. 19 schematically shows a different thread of control of aprocessor operating in a plurality of modes;

[0074]FIG. 20 schematically shows a further thread of control of aprocessor operating in a plurality of modes, wherein interrupts areenabled in monitor mode;

[0075] FIGS. 21 to 23 illustrate a view of different processing modesand scenario for switching between secure and non-secure domains inaccordance with another example embodiment(s);

[0076]FIG. 24 schematically illustrates the concept of adding a secureprocessing option to a traditional ARM core;

[0077]FIG. 25 schematically illustrates a processor having a secure andnon-secure domain and reset;

[0078]FIG. 26 schematically illustrates the delivering of processingrequests to a suspended operating system using a software fakedinterrupt;

[0079]FIG. 27 schematically illustrates another example of thedelivering of a processing request to a suspended operating system via asoftware faked interrupt;

[0080]FIG. 28 is a flow diagram schematically illustrating processingperformed upon receipt of a software faked interrupt of the typegenerated in FIGS. 26 and 27;

[0081]FIGS. 29 and 30 schematically illustrate task following by asecure operating system to track possible task switches made by anon-secure operating system;

[0082]FIG. 31 is a flow diagram schematically illustrating theprocessing performed upon receipt of a call at the secure operatingsystem of FIGS. 29 and 30;

[0083]FIG. 32 is a diagram schematically illustrating the problem ofinterrupt priority inversion which may occur in a system having multipleoperating systems where different interrupts may be handled by differentoperating systems;

[0084]FIG. 33 is a diagram schematically illustrating the use of stubinterrupt handlers to avoid the problem illustrated in FIG. 32; and

[0085]FIG. 34 schematically illustrates how different types andpriorities of interrupts may be handled depending upon whether or notthey can be interrupted by an interrupt which will be serviced using adifferent operating system;

[0086]FIG. 35 illustrates how the processor configuration data isoverridden with monitor mode specific processor configuration data whenthe processor is operating in monitor mode;

[0087]FIG. 36 is a flow diagram illustrating how the processorconfiguration data is switched when transitioning between the securedomain and the non-secure domain in accordance with one embodiment tothe present invention;

[0088]FIG. 37 is a diagram illustrating the memory management logic usedin one embodiment of the present invention to control access to memory;

[0089]FIG. 38 is a block diagram illustrating the memory managementlogic of a second embodiment of the present invention used to controlaccess to memory;

[0090]FIG. 39 is a flow diagram illustrating the process performed inone embodiment of the present invention within the memory managementlogic to process a memory access request that specifies a virtualaddress;

[0091]FIG. 40 is a flow diagram illustrating the process performed inone embodiment of the present invention within the memory managementlogic to process a memory access request that specifies a physicaladdress;

[0092]FIG. 41 schematically illustrates how the partition checker ofpreferred embodiments is operable to prevent access to a physicaladdress within secure memory when the device issuing the memory accessrequest is operating in a non-secure mode;

[0093]FIG. 42 is a diagram illustrating the use of both a non-securepage table and a secure page table in preferred embodiments of thepresent invention;

[0094]FIG. 43 is a diagram illustrating two forms of flag used withinthe main translation lookaside buffer (TLB) of preferred embodiments;

[0095]FIG. 44 illustrates how memory may be partitioned after a bootstage in one embodiment of the present invention;

[0096]FIG. 45 illustrates the mapping of the non-secure memory by thememory management unit following the performance of the boot partitionin accordance with an embodiment of the present invention;

[0097]FIG. 46 illustrates how the rights of a part of memory can bealtered to allow a secure application to share memory with a non-secureapplication in accordance with an embodiment of the present invention;

[0098]FIG. 47 illustrates how devices may be connected to the externalbus of the data processing apparatus in accordance with one embodimentof the present invention;

[0099]FIG. 48 is a block diagram illustrating how devices may be coupledto the external bus in accordance with the second embodiment of thepresent invention;

[0100]FIG. 49 illustrates the arrangement of physical memory inembodiments where a single set of page tables is used;

[0101]FIG. 50A illustrates an arrangement in which two MMUs are used toperform virtual to physical address translation via an intermediateaddress;

[0102]FIG. 50B illustrates an alternative arrangement in which two MMUsare used to perform virtual to physical address translation via anintermediate address;

[0103]FIG. 51 illustrates, by way of example, the correspondence betweenphysical address space and intermediate address space for both thesecure domain and the non-secure domain;

[0104]FIG. 52 illustrates the swapping of memory regions between secureand non-secure domains through manipulation of the page tablesassociated with the second MMU;

[0105]FIG. 53 is an embodiment illustrating an implementation using asingle MMU, and where a miss in the main TLB causes an exception to beinvoked to determine the virtual to physical address translation;

[0106]FIG. 54 is a flow diagram illustrating the process performed bythe processor core in order to action an exception issued uponoccurrence of a miss in the main TLB of the MMU of FIG. 53;

[0107]FIG. 55 is a block diagram illustrating components provided withina data processing apparatus of one embodiment, in which the cache isprovided with information as to whether the data stored in individualcache lines is secure data or non-secure data;

[0108]FIG. 56 illustrates the construction of the memory management unitillustrated in FIG. 55;

[0109]FIG. 57 is a flow diagram illustrating the processing performedwithin the data processing apparatus of FIG. 55 to process a non-securememory access request;

[0110]FIG. 58 is a flow diagram illustrating the processing performedwithin the data processing apparatus of FIG. 55 in order to process asecure memory access request;

[0111]FIG. 59 schematically shows possible granularity of monitoringfunctions for different modes and applications running on a processor;

[0112]FIG. 60 shows possible ways of initiating different monitoringfunctions;

[0113]FIG. 61 shows a table of control values for controllingavailability of different monitoring functions;

[0114]FIG. 62 shows a positive-edge triggered FLIP-FLOP view;

[0115]FIG. 63 a scan chain cell;

[0116]FIG. 64 shows a plurality of scan chain cells in a scan chain;

[0117]FIG. 65 shows a debug TAP controller;

[0118]FIG. 66A shows a debug TAP controller with a JADI input;

[0119]FIG. 66B shows a scan chain cell with a bypass register;

[0120]FIG. 67 schematically illustrates a processor comprising a core,scan chains and a Debug Status and Control Register;

[0121]FIG. 68 schematically illustrates the factors controlling debug ortrace initialisation;

[0122]FIGS. 69A and 69B show a summary of debug granularity;

[0123]FIG. 70 schematically illustrates the granularity of debug whileit is running; and

[0124]FIGS. 71A and 71B show monitor debug when debug is enabled insecure world and when it is not enabled respectively.

[0125]FIG. 1 is a block diagram illustrating a data processing apparatusin accordance with preferred embodiments of the present invention. Thedata processing apparatus incorporates a processor core 10 within whichis provided an arithmetic logic unit (ALU) 16 arranged to executesequences of instructions. Data required by the ALU 16 is stored withina register bank 14. The core 10 is provided with various monitoringfunctions to enable diagnostic data to be captured indicative of theactivities of the processor core. As an example, an Embedded TraceModule (ETM) 22 is provided for producing a real time trace of certainactivities of the processor core in dependence on the contents ofcertain control registers 26 within the ETM 22 defining which activitiesare to be traced. The trace signals are typically output to a tracebuffer from where they can subsequently be analysed. A vectoredinterrupt controller 21 is provided for managing the servicing of aplurality of interrupts which may be raised by various peripherals (notillustrated).

[0126] Further, as shown in FIG. 1, another monitoring functionalitythat can be provided within the core 10 is a debug function, a debuggingapplication external to the data processing apparatus being able tocommunicate with the core 10 via a Joint Test Access Group (JTAG)controller 18 which is coupled to one or more scan chains 12.Information about the status of various parts of the processor core 10can be output via the scan chains 12 and the JTAG controller 18 to theexternal debugging application. An In Circuit Emulator (ICE) 20 is usedto store within registers 24 conditions identifying when the debugfunctions should be started and stopped, and hence for example will beused to store breakpoints, watchpoints, etc.

[0127] The core 10 is coupled to a system bus 40 via memory managementlogic 30 which is arranged to manage memory access requests issued bythe core 10 for access to locations in memory of the data processingapparatus. Certain parts of the memory may be embodied by memory unitsconnected directly to the system bus 40, for example the Tightly CoupledMemory (TCM) 36, and the cache 38 illustrated in FIG. 1. Additionaldevices may also be provided for accessing such memory, for example aDirect Memory Access (DMA) controller 32. Typically, various controlregisters 34 will be provided for defining certain control parameters ofthe various elements of the chip, these control registers also beingreferred to herein as coprocessor 15 (CP15) registers.

[0128] The chip containing the core 10 may be coupled to an external bus70 (for example a bus operating in accordance with the “AdvancedMicrocontroller Bus Architecture” (AMBA) specification developed by ARMLimited) via an external bus interface 42, and various devices may beconnected to the external bus 70. These devices may include masterdevices such as a digital signal processor (DSP) 50, or a direct memoryaccess (DMA) controller 52, as well as various slave devices such as theboot ROM 44, the screen driver 46, the external memory 56, aninput/output (I/O) interface 60 or a key storage unit 64. These variousslave devices illustrated in FIG. 1 can all be considered asincorporating parts of the overall memory of the data processingapparatus. For example, the boot ROM 44 will form part of theaddressable memory of the data processing apparatus, as will theexternal memory 56. Further, devices such as the screen driver 46, I/Ointerface 60 and key storage unit 64 will all include internal storageelements such as registers or buffers 48, 62, 66, respectively, whichare all independently addressable as part of the overall memory of thedata processing apparatus. As will be discussed in more detail later, apart of the memory, e.g. part of external memory 56, will be used tostore one or more page tables 58 defining information relevant tocontrol of memory accesses.

[0129] As will be appreciated by those skilled in the art, the externalbus 70 will typically be provided with arbiter and decoder logic 54, thearbiter being used to arbitrate between multiple memory access requestsissued by multiple master devices, for example the core 10, the DMA 32,the DSP 50, the DMA 52, etc, whilst the decoder will be used todetermine which slave device on the external bus should handle anyparticular memory access request.

[0130] Whilst in some embodiments, the external bus may be providedexternally to the chip containing the core 10, in other embodiments theexternal bus will be provided on-chip with the core 10. This has thebenefit that secure data on the external bus is easier to keep securethan when the external bus is off-chip; when the external bus isoff-chip, data encryption techniques may be used to increase thesecurity of secure data.

[0131]FIG. 2 schematically illustrates various programs running on aprocessing system having a secure domain and a non-secure domain. Thesystem is provided with a monitor program 72 which executes at leastpartially in a monitor mode. In this example embodiment security statusflag is write accessible only within the monitor mode and may be writtenby the monitor program 72. The monitor program 72 is responsible formanaging all changes between the secure domain and the non-secure domainin either direction. From a view external to the core the monitor modeis always secure and the monitor program is in secure memory.

[0132] Within the non-secure domain there is provided a non-secureoperating system 74 and a plurality of non-secure application programs76, 78 which execute. in cooperation with the non-secure operatingsystem 74. In the secure domain, a secure kernel program 80 is provided.The secure kernel program 80 can be considered to form a secureoperating system. Typically such a secure kernel program 80 will bedesigned to provide only those functions which are essential toprocessing activities which must be provided in the secure domain suchthat the secure kernel 80 can be as small and simple as possible sincethis will tend to make it more secure. A plurality of secureapplications 82, 84 are illustrated as executing in combination with thesecure kernel 80.

[0133]FIG. 3 illustrates a matrix of processing modes associated withdifferent security domains. In this particular example the processingmodes are symmetrical with respect to the security domain andaccordingly Mode 1 and Mode 2 exist in both secure and non-secure forms.

[0134] The monitor mode has the highest level of security access in thesystem and in this example embodiment is the only mode entitled toswitch the system between the non-secure domain and the secure domain ineither direction. Thus, all domain switches take place via a switch tothe monitor mode and the execution of the monitor program 72 within themonitor mode.

[0135]FIG. 4 schematically illustrates another set of non-secure domainprocessing modes 1, 2, 3, 4 and secure domain processing modes a, b, c.In contrast to the symmetric arrangement of FIG. 3, FIG. 4 shows thatsome of the processing modes may not be present in one or other of thesecurity domains. The monitor mode 86 is again illustrated as straddlingthe non-secure domain and the secure domain. The monitor mode 86 can beconsidered a secure processing mode, since the secure status flag may bechanged in this mode and monitor program 72 in the monitor mode has theability to itself set the security status flag it effectively providesthe ultimate level of security within the system as a whole.

[0136]FIG. 5 schematically illustrates another arrangement of processingmodes with respect to security domains. In this arrangement both secureand non-secure domains are identified as well as a further domain. Thisfurther domain may be such that it is isolated from other parts of asystem in a way that it does not need to interact with either of thesecure domain or non-secure domain illustrated and as such the issue ofto which of these it belongs to is not relevant.

[0137] As will be appreciated a processing system, such as amicroprocessor is normally provided with a register bank 88 in whichoperand values may be stored. FIG. 6 illustrates a programmer's modelview of an example register bank with dedicated registers being providedfor certain of the register numbers in certain of the processing modes.More particularly, the example of FIG. 6 is an extension of the knownARM register bank (e.g. as provided in ARM7 processors of ARM Limited,Cambridge, England) which is provided with a dedicated saved programstatus register, a dedicated stack pointer register and a dedicated linkregister R14 for each processing mode, but in this case extended by theprovision of a monitor mode. As illustrated in FIG. 6, the fastinterrupt mode has additional dedicated registers provided such thatupon entry of the fast interrupt mode there is no need to save and thenrestore register contents from other modes. The monitor mode may inalternative embodiments also be provided with dedicated furtherregisters in a similar manner to the fast interrupt mode so as to speedup processing of a security domain switch and reduce system latencyassociated with such switches.

[0138]FIG. 7 schematically illustrates another embodiment in which theregister bank 88 is provided in the form of two complete and separateregister banks that are respectively used in the secure domain and thenon-secure domain. This is one way in which secure data stored withinregisters operable in the secure domain can be prevented from becomingaccessible when a switch is made to the non-secure domain. However, thisarrangement hinders the possibility of passing data from the non-securedomain to the secure domain as may be permitted and desirable by usingthe fast and efficient mechanism of placing it in a register which isaccessible in both the non-secure domain and the secure domain.

[0139] An important advantage of having secure register bank is to avoidthe need for flushing the contents of registers before switching fromone world to the other. If latency is not a critical issue, a simplerhardware system with no duplicated registers for the secure domain worldmay be used, e.g. FIG. 6. The monitor mode is responsible switching fromone domain to the other. Restoring context, saving previous context, aswell as flushing registers is performed by a monitor program at leastpartially executing in monitor mode. The system behaves thus like avirtualisation model. This type of embodiment is discussed furtherbelow. Reference should be made to, for example, the programmer's modelof the ARM7 upon which the security features described herein build.

[0140] Processor Modes

[0141] Instead of duplicating modes in secure world, the same modessupport both secure and non-secure domains (see FIG. 8). Monitor mode isaware of the current status of the core, either secure or non-secure(e.g. as read from an S bit stored is a coprocessor configurationregister).

[0142] In the FIG. 8, whenever an SMI (Software Monitor Interruptinstruction) occurs, the core enters monitor mode to switch properlyfrom one world to the other.

[0143] With reference to FIG. 9 in which SMIs are permitted from usermode:

[0144] 1. The scheduler launches thread 1

[0145] 2. Thread 1 needs to perform a secure function=>SMI secure call,the core enters monitor mode. Under hardware control the current PC andCPSR (current processor status register) are stored in R14_mon andSPSR_mon (saved processor status register for the monitor mode) andIRQ/FIQ interrupts are disabled.

[0146] 3. The monitor program does the following tasks:

[0147] The S bit is set (the secure status flag).

[0148] Saves at least R14_mon and SPSR_mon in a stack so that non-securecontext cannot be lost if an exception occurs whilst the secureapplication is running.

[0149] Checks there is a new thread to launch: secure thread 1. Amechanism (via thread ID table in some example embodiments) indicatesthat thread 1 is active in the secure world.

[0150] IRQ/FIQ interrupts are re-enabled. A secure application can thenstart in secure user mode.

[0151] 4. Secure thread 1 runs until it finishes, then branches (SMI)onto the ‘return from secure’ function of the monitor program mode(IRQ/FIQ interrupts are then disabled when the core enters monitor mode)

[0152] 5. The ‘return from secure’ function does the following tasks:

[0153] indicates that secure thread 1 is finished (e.g., in the case ofa thread ID table, remove thread 1 from the table).

[0154] Restore from stack non-secure context and flush requiredregisters, so that no secure data can be read once return has been madeto the non-secure domain.

[0155] Then branches back to the non-secure domain with a SUBSinstruction (this restores the program counter to the correct point andupdates the status flags), restoring the PC (from restored R14_mon) andCPSR (from SPSR_mon). So, the return point in the non-secure domain isthe instruction following the previously executed SMI in thread 1.

[0156] 6. Thread 1 executes until the end, then gives the hand back tothe scheduler.

[0157] Some of the above functionality may be split between the monitorprogram and the secure operating system depending upon the particularembodiment.

[0158] In other embodiments it may be desired not to allow SMIs to occurin user modes.

[0159] Secure World Entry

Reset

[0160] When a hardware reset occurs, the MMU is disabled and the ARMcore (processor) branches to secure supervisor mode with the S bit set.Once the secure boot is terminated an SMI to go to monitor mode may beexecuted and the monitor can switch to the OS in non-secure world(non-secure svc mode) if desired. If it is desired to use a legacy OSthis can simply boot in secure supervisor mode and ignore the securestate.

SMI INSTRUCTION

[0161] This instruction (a mode switching software interruptinstruction) can be called from any non-secure modes in the non-securedomain (as previously mentioned it may be desired to restrict SMIs toprivileged modes), but the target entry point determined by theassociated vector is always fixed and within monitor mode. Its up to theSMI handler to branch to the proper secure function that must be run(e.g. controlled by an operand passed with the instruction).

[0162] Passing parameters from non-secure world to secure world can beperformed using the shared registers of the register bank within a FIG.6 type register bank.

[0163] When a SMI occurs in non-secure world, the ARM core may do thefollowing actions in hardware:

[0164] Branch to SMI vector (in secure memory access is allowed sinceyou will now be in monitor mode) into monitor mode

[0165] Save PC into R14_mon and CPSR into SPSR_mon

[0166] Set the S bit using the monitor program

[0167] Start to execute secure exception handler in monitor mode(restore/save context in case of multi-threads)

[0168] Branch to secure user mode (or another mode, like svc mode) toexecute the appropriate function

[0169] IRQ and FIQ are disabled while the core is in monitor mode(latency is increased)

Secure World Exit

[0170] There are two possibilities to exit secure world:

[0171] The secure function is finished and we return into previousnon-secure mode that had called this function.

[0172] The secure function is interrupted by a non-secure exception(e.g. IRQ/FIQ/SMI).

Normal End of Secure Function

[0173] The secure function terminates normally and we need to resume anapplication in the non-secure world at the instruction just after theSMI. In the secure user mode, a ‘SMI’ instruction is performed to returnto monitor mode with the appropriate parameters corresponding to a‘return from secure world’ routine. At this stage, the registers areflushed to avoid leakage of data between non-secure and secure worlds,then non-secure context general purpose registers are restored andnon-secure banked registers are updated with the value they had innon-secure world. R14_mon and SPSR_mon thus get the appropriate valuesto resume the non-secure application after the SMI, by executing a ‘MOVSPC, R14’ instruction.

Exit of Secure Function Due to a Non-Secure Exception

[0174] In this case, the secure function is not finished and the securecontext must be saved before going into the non-secure exceptionhandler, whatever the interrupts are that need to be handled.

Secure Interrupts

[0175] There are several possibilities for secure interrupts.

[0176] Two possible solutions are proposed which depend on:

[0177] What kind of interrupt it is (secure or non-secure)

[0178] What mode the core is in when the IRQ occurs (either in secure orin non-secure world)

Solution One

[0179] In this solution, two distinct pins are required to supportsecure and non-secure interrupts.

[0180] While in Non Secure world, if

[0181] an IRQ occurs, the core goes to IRQ mode to handle this interruptas in ARM cores such as the ARM7

[0182] a SIRQ occurs, the core goes to monitor mode to save non-securecontext and then to a secure IRQ handler to deal with the secureinterrupt.

[0183] While in Secure world, if

[0184] an SIRQ occurs, the core goes to the secure IRQ handler. The coredoes not leave the secure world

[0185] an IRQ occurs, the core goes to monitor mode where secure contextis saved, then to a non-secure IRQ handler to deal with this non-secureinterrupt.

[0186] In other words, when an interrupt that does not belong to thecurrent world occurs, the core goes directly to monitor mode, otherwiseit stays in the current world (see FIG. 10).

IRO Occurring in Secure World

[0187] See FIG. 11A:

[0188] 1. The scheduler launches thread 1.

[0189] 2. Thread 1 needs to perform a secure function=>SMI secure call,the core enters monitor mode. Current PC and CPSR are stored in R14_monand SPSR_mon, IRQ/FIQ are disabled.

[0190] 3. The monitor handler (program) does the following tasks:

[0191] The S bit is set.

[0192] Saves at least R14_mon and SPSR_mon in a stack (and possiblyother registers are also pushed) so that non-secure context cannot belost if an exception occurs whilst the secure application is running.

[0193] Checks there is a new thread to launch: secure thread 1. Amechanism (via thread ID table) indicates that thread 1 is active in thesecure world.

[0194] Secure application can then start in the secure user mode.IRQ/FIQ are then re-enabled.

[0195] 4. An IRQ occurs while secure thread 1 is running. The core jumpsdirectly to monitor mode (specific vector) and stores current PC inR14_mon and CPSR in SPSR_mon in monitor mode, (IRQ/FIQ are thendisabled).

[0196] 5. Secure context must be saved, previous non-secure context isrestored. The monitor handler may be to IRQ mode to updateR14_irq/SPSR_irq with appropriate values and then passes control to anon-secure IRQ handler.

[0197] 6. The IRQ handler services the IRQ, then gives control back tothread 1 in the non-secure world. By restoring SPRS_irq and R14_irq intothe CPSR and PC, thread 1 is now pointing onto the SMI instruction thathas been interrupted.

[0198] 7. The SMI instruction is re-executed (same instruction as 2).

[0199] 8. The monitor handler sees this thread has previously beeninterrupted, and restores the thread 1 context. It then branches tosecure thread 1 in user mode, pointing onto the instruction that hasbeen interrupted.

[0200] 9. Secure thread 1 runs until it finishes, then branches onto the‘return from secure’ function in monitor mode (dedicated SMI).

[0201] 10. The ‘return from secure’ function does the following tasks:

[0202] indicates that secure thread 1 is finished (i.e., in the case ofa thread ID table, remove thread 1 from the table).

[0203] restore from stack non-secure context and flush requiredregisters, so that no secure data can be read once a return is made tonon-secure world.

[0204] branches back to the non-secure world with a SUBS instruction,restoring the PC (from restored R14_mon) and CPSR (from SPSR_mon). So,the return point in the non-secure world should be the instructionfollowing the previously executed SMI in thread 1.

[0205] 11. Thread 1 executes until the end, then gives control back tothe scheduler.

SIRO Occurring in Non-Secure World

[0206] See FIG. 11B:

[0207] 1. The schedule launches thread 1

[0208] 2. A SIRQ occurs while secure thread 1 is running. The core jumpsdirectly to monitor mode (specific vector) and stores current PC inR14_mon and CPSR in SPSR_mon in monitor mode, IRQ/FIQ are then disabled.

[0209] 3. Non-Secure context must be saved, then the core goes to asecure IRQ handler.

[0210] 4. The IRQ handler services the SIRQ, then gives control back tothe monitor mode handler using an SMI with appropriate parameters.

[0211] 5. The monitor handler restores non-secure context so that a SUBSinstruction makes the core return to the non-secure world and resumesthe interrupted thread 1.

[0212] 6. Thread 1 executes until the end, then gives the hand back tothe scheduler.

[0213] The mechanism of FIG. 11A has the advantage of providing adeterministic way to enter secure world. However, there are someproblems associated with interrupt priority: e.g. while a SIRQ isrunning in secure interrupt handler, a non-secure IRQ with higherpriority may occur. Once the non-secure IRQ is finished, there is a needto recreate the SIRQ event so that the core can resume the secureinterrupt.

Solution Two

[0214] In this mechanism (See FIG. 12) two distinct pins, or only one,may support secure and non-secure interrupts. Having two pins reducesinterrupt latency.

[0215] While in Non Secure world, if

[0216] an IRQ occurs, the core goes to IRQ mode to handle this interruptlike in ARM7 systems

[0217] a SIRQ occurs, the core goes to an IRQ handler where an SMIinstruction will make the core branch to monitor mode to save non-securecontext and then to a secure IRQ handler to deal with the secureinterrupt.

[0218] While in a Secure world, if

[0219] a SIRQ occurs, the core goes to the secure IRQ handler. The coredoes not leave the secure world

[0220] an IRQ occurs, the core goes to the secure IRQ handler where anSMI instruction will make the core branch to monitor mode (where securecontext is saved), then to a non-secure IRQ handler to deal with thisnon-secure interrupt.

IRO Occurring In Secure World

[0221] See FIG. 13A:

[0222] 1. The schedule launches thread 1.

[0223] 2. Thread 1 needs to perform a secure function=>SMI secure call,the core enters monitor mode. Current PC and CPSR are stored in R14_monand SPSR_mon, IRQ/FIQ are disabled.

[0224] 3. The monitor handler does the following tasks:

[0225] The S bit is set.

[0226] Saves at least R14_mon and SPSR_mon in a stack (eventually otherregisters) so that non-secure context cannot be lost if an exceptionoccurs whilst the secure application is running.

[0227] Checks there is a new thread to launch: secure thread 1. Amechanism (via thread ID table) indicates that thread 1 is active in thesecure world.

[0228] Secure application can then start in the secure user mode.IRQIFIQ are re-enabled.

[0229] 4. An IRQ occurs while secure thread 1 is running. The core jumpsdirectly to secure IRQ mode.

[0230] 5. The core stores current PC in R14_irq and CPSR in SPSR_irq.The IRQ handler detects this is a non-secure interrupt and performs aSMI to enter monitor mode with appropriate parameters.

[0231] 6. Secure context must be saved, previous non-secure context isrestored. The monitor handler knows where the SMI came from by readingthe CPSR. It can also go to IRQ mode to read R14_irq/SPSR_irq to saveproperly secure context. It can also save in these same registers thenon-secure context that must be restored once the IRQ routine will befinished.

[0232] 7. The IRQ handler services the IRQ, then gives control back tothread 1 in the non-secure world. By restoring SPRS_irq and R14_irq intothe CPSR and PC, the core is now pointing onto the SMI instruction thathas been interrupted.

[0233] 8. The SMI instruction is re-executed (same instruction as 2).

[0234] 9. The monitor handler sees this thread has previously beeninterrupted, and restores the thread 1 context. It then branches tosecure thread 1 in user mode, pointing to the instruction that has beeninterrupted.

[0235] 10. Secure thread 1 runs until it finishes, then branches ontothe ‘return from secure’; function in monitor mode (dedicated SMI).

[0236] 11. The ‘return from secure’ function does the following tasks:

[0237] indicates that secure thread 1 is finished (i.e., in the case ofa thread ID table, remove thread 1 from the table).

[0238] restores from stack non-secure context and flushes requiredregisters, so that no secure information can be read once we return innon-secure world.

[0239] branches back to the non-secure world with a SUBS instruction,restoring the PC (from restored R14_mon) and CPSR (from SPSR_mon). Thereturn point in the non-secure world should be the instruction followingthe previously executed SMI in thread 1.

[0240] 12. Thread 1 executes until the end, then gives the hand back tothe scheduler.

SIRO Occurring in Non-Secure World

[0241] See FIG. 13B:

[0242] 1. The schedule launches thread 1.

[0243] 2. A SIRQ occurs while secure thread 1 is running.

[0244] 3. The core jumps directly irq mode and stores current PC inR14_irq and CPSR in SPSR_irq. IRQ is then disabled. The IRQ handlerdetects this is a SIRQ and performs a SMI instruction with appropriateparameters.

[0245] 4. Once in monitor mode, non-secure context must be saved, thenthe core goes to a secure IRQ handler.

[0246] 5. The secure IRQ handler services the SIRQ service routine, thengives control back to monitor with SMI with appropriate parameters.

[0247] 6. The monitor handler restores non-secure context so that a SUBSinstruction makes the core returns to non-secure world and resumes theinterrupted IRQ handler.

[0248] 7. The IRQ handler may then return to the non-secure thread byperforming a SUBS.

[0249] 8. Thread 1 executes until the end, then gives control back tothe scheduler.

[0250] With the mechanism of FIG. 12, there is no need to recreate theSIRQ event in the case of nested interrupts, but there is no guaranteethat secure interrupts will be performed.

[0251] Exception Vectors

[0252] At least two physical vector tables are kept (although from avirtual address point of view they may appear. as a single vectortable), one for the non-secure world in non-secure memory, the one forthe secure world in secure memory (not accessible from non-secureworld). The different virtual to physical memory mappings used in thesecure and non-secure worlds effectively allow the same virtual memoryaddresses to access different vector tables stored in physical memory.The monitor mode may always use flat memory mapping to provide a thirdvector table in physical memory.

[0253] If the interrupts follow the FIG. 12 mechanism, there would bethe following vectors shown in FIG. 14 for each table. This vector setis duplicated in both secure and non-secure memory. Exception VectorOffset Corresponding Mode Reset 0x00 Supervisor Mode (S bit set) Undef0x04 Monitor mode/Undef mode SWI 0x08 Supervisor mode/Monitor modePrefetch Abort 0x0C Abort mode/Monitor mode Data Abort 0x10 Abortmode/Monitor Mode IRQ/SIRQ 0x18 IRQ mode FIQ 0x1X FIQ mode SMI 0x20Undef mode/Monitor mode

[0254] NB. The Reset entry is only in the secure vector table. When aReset is performed in non secure world, the core hardware forces entryof supervisor mode and setting of the S bit so that the Reset vector canbe accessed in secure memory.

[0255]FIG. 15 illustrates three exception vector tables respectivelyapplicable to a secure mode, a non-secure mode and the monitor mode.These exception vector tables may be programmed with exception vectorsin order to match the requirements and characteristics of the secure andnon-secure operating systems. Each of the exception vector tables mayhave an associated vector table base address register within CP15storing a base address pointing to that table within memory. When anexception occurs the hardware will reference the vector table baseaddress register corresponding to the current state of the system todetermine the base address of the vector table to be used.Alternatively, the different virtual to physical memory mappings appliedin the different modes may be used to separate the three differentvector table stored at different physical memory addresses. Asillustrated in FIG. 16, an exception trap mask register is provided in asystem (configuration controlling) coprocessor (CP15) associated withthe processor core. This exception trap mask register provides flagsassociated with respective exception types. These flags indicate whetherthe hardware should operate to direct processing to either the vectorfor the exception concerned within its current domain or should force aswitch to the monitor mode (which is a type of secure mode) and thenfollow the vector in the monitor mode vector table. The exception trapmask register (exception control register) is only writable from themonitor mode. It may be that read access is also prevented to theexception trap mask register when in a non-secure mode. It will be seenthat the exception trap mask register of FIG. 16 does not include a flagfor the reset vector as the system is configured to always force this tojump to the reset vector in the secure supervisor mode as specified inthe secure vector table in order to ensure a secure boot and backwardscompatibility. It will be seen that in FIG. 15, for the sake ofcompleteness, reset vectors have been shown in the vector tables otherthan the secure supervisor mode secure vector table.

[0256]FIG. 16 also illustrates that the flags for the differentexception types within the exception trap mask register areprogrammable, such as by the monitor program during secure boot.Alternatively, some or all of the flags may in certain implementationsbe provided by physical input signals, e.g. the secure interrupt flagSIRQ may be hardwired to always force monitor mode entry and executionof the corresponding monitor mode secure interrupt request vector when asecure interrupt signal is received. FIG. 16 illustrates only thatportion of the exception trap register concerned with non-secure domainexceptions, a similar set of programmable bits will be provided forsecure domain exceptions.

[0257] Whilst it will be understood from the above that at one level thehardware acts to either force an interrupt to be serviced by the currentdomain exception handler or the monitor mode exception handler dependingupon the exception control register flags, this is only the first levelof control that is applied. As an example, it is possible for anexception to occur in the secure mode, the secure mode exception vectorto be followed to the secure mode exception handler, but this securemode exception handler then decide that the exception is of a naturethat it is better dealt with by the non-secure exception handler andaccordingly utilise an SMI instruction to switch to the non-secure modeand invoke the non-secure exception handler. The converse is alsopossible where the hardware might act to initiate the non-secureexception handler, but this then execute instructions which directprocessing to the secure exception handler or the monitor mode exceptionhandler.

[0258]FIG. 17 is a flow diagram schematically illustrating the operationof the system so as to support another possible type of switchingrequest associated with a new type of exception. At step 98 the hardwaredetects any instruction which is attempting to change to monitor mode asindicate in a current program status register (CPSR). When such anattempt is detected, then a new type of exception is triggered, thisbeing referred to herein as a CPSR violation exception. The generationof this CPSR violation exception at step 100 results in reference to anappropriate exception vector within the monitor mode and the monitorprogram is run at step 102 to handle the CPSR violation exception.

[0259] It will be appreciated that the mechanisms for initiating aswitch between secure domain and non-secure domain discussed in relationto FIG. 17 may be provided in addition to support for the SMIinstruction previously discussed. This exception mechanism may beprovided to respond to unauthorised attempts to switch mode as allauthorised attempts should be made via an SMI instruction.Alternatively, such a mechanism may be legitimate ways to switch betweenthe secure domain and the non-secure domain or may be provided in orderto give backwards compatibility with existing code which, for example,might seek to clear the processing status register as part of its normaloperation even though it was not truly trying to make an unauthorisedattempt to switch between the secure domain and the non-secure domain.

[0260] As described above, in general interrupts are disabled when theprocessor is operating in monitor mode. This is done to increase thesecurity of the system. When an interrupt occurs the state of theprocessor at that moment is stored in interrupt exception registers sothat on completion of the interrupt function the processing of theinterrupted function can be resumed at the interrupt point. If thisprocess were allowed in monitor mode it could reduce the security of themonitor mode, giving a possible secure data leakage path. For thisreason interrupts are generally disabled in monitor mode. However, oneconsequence of disabling interrupts during monitor mode is thatinterrupt latency is increased.

[0261] It would be possible to allow interrupts in monitor mode if thestate of the processor executing the function was not stored. This canonly be done if following an interrupt the function is not resumed.Thus, the problem of interrupt latency in monitor mode may be addressedby allowing interrupts in monitor mode only of functions that can besafely restarted. In this case, following an interrupt in monitor mode,the data relating to the processing of the function is not stored but isthrown away and the processor is instructed to start processing of thefunction from its beginning once the interrupt has finished. In theabove example this is a simple thing to do as the processor simplyreturns to the point at which it switched to monitor mode. It should benoted that restarting a function is only possible for certain functionsthat can be restarted and still produce repeatable results. If thefunction has changed a state of the processor such that if it wererestarted it would produce a different result then it is not a good ideato restart the function. For this reason, only those functions that aresafely restartable can be interrupted in monitor mode, for otherfunctions the interrupts are disabled.

[0262]FIG. 18 illustrates a way of dealing with an interrupt occurringin monitor mode according to an embodiment of the present invention. AnSMI occurs during processing of task A in a non-secure mode and thisswitches the processor to monitor mode. The SMI instruction makes thecore enter the Monitor mode through a dedicated non-secure SMI vector.The current state of the PC is saved, the s bit is set and interruptsare disabled. Generally, LR_mon and SPSR_mon are used to save the PC andCPSR of the non secure mode.

[0263] A function, function C is then initiated in monitor mode. Thefirst thing function C does is to enable the interrupts, function C isthen processed. If an interrupt occurs during the processing of functionC, the interrupts are not disabled so the interrupt is accepted andperformed. However, the monitor mode indicator indicates to theprocessor that following an interrupt, the function is not to beresumed, but rather restarted. Alternatively, this may be indicated tothe processor by a separate control parameter. Thus, following aninterrupt the interrupt exception vectors are updated with the values ofLR_mon and SPSR_mon and the current state of the processor is notstored.

[0264] As is shown in FIG. 18 following completion of the interrupttask, task B, the processor reads the address of the SMI instructionwhich has been copied to the interrupt register and performs an SMI andstarts to process function C again.

[0265] The above process only works if function C is restartable, thatis to say if restarting process C will result in repeatable processingsteps. This will not be the case if function C has changed any of thestates of the processor such as the stack pointer that may affect itsfuture processing. A function that is repeatable in this way is said tohave idempotence. One way of dealing with the problem of a function nothaving idempotence is to rearrange the code defining the function insuch a way that the first portion of the code has idempotence and onceit is no longer possible to arrange the code to have idempotenceinterrupts are disabled. For example, if code C involves writing to thestack, it may be possible to do so without updating the stack pointer atleast at first. Once it is decided that the code can no longer feasiblybe safely restarted, then the code for function C can instruct theprocessor to disable interrupts and then it can update the stack pointerto the correct position. This is shown in FIG. 18 where interrupts aredisabled a certain way through the processing of function C.

[0266]FIG. 19 illustrates a slightly different example. In this example,a certain way through the processing of task C, a further controlparameter is set. This indicates that the following portion of task C isnot strictly idempotent, but can be safely restarted provided that afix-up routine is run first. This fix-up routine acts to restore a stateof the processor to how it was at the start of task C, such that task Ccan be safely restarted and produce the same processor state at the endof the task as it would have done had it not been interrupted. In someembodiments at the point that the further control parameter is setinterrupts may be disabled for a short while while some states of theprocessor are amended such as the stack pointer being updated. Thisallows the processor to be restored to an idempotent state later.

[0267] When an interrupt occurs after the further control parameter hasbeen set, then there are two possible ways to proceed. Either the fix-uproutine can be performed immediately (at F1) and then the interrupt canbe processed, or the interrupt can be processed immediately andfollowing completion of the interrupt, the SMI is executed and thenprior to restarting task C the fix-up routine is performed (at F2). Ascan be seen, in both of these embodiments the fix-up routine isperformed in monitor mode, and thus execution in the non-secure domain,which is not aware of the secure domain or of the monitor mode is notaffected.

[0268] As can be seen from FIG. 19, a first portion of code C hasidempotence and can be restarted following an interrupt. A secondportion is restartable provided a fix-up routine is run first, and thisis indicated by setting a “further” control parameter, and a finalportion of the code cannot be restarted and thus, interrupts aredisabled before this code is processed.

[0269]FIG. 20 illustrates an alternative example, in this case, which isdifferent to other embodiments, interrupts are enabled during themonitor mode. Functions running in the monitor mode then act to disableinterrupts as soon as they are no longer safely restartable. This isonly possible if all functions interrupted in monitor mode are restartedrather than resumed.

[0270] There are several ways that it can be ensured that all functionsrunning in a certain mode are restarted rather than resumed wheninterrupted. One way is by adding a new processor state in whichinterrupts save the address of the start of the instruction sequencerather than the address of the interrupted instruction. In this casemonitor mode would then always be run in this state. An alternative wayis by preloading the address of the start of a function to the interruptexception register at the start of each function and disablingsubsequent writing of the state of the processor following interrupt tointerrupt exception registers.

[0271] In the embodiment illustrated in FIG. 20 restarting of thefunctions may be done immediately following termination of the interruptfunction or it may be done following a fix-up routine, if that isrequired to make the function safely restartable.

[0272] Although the above described way of dealing with interruptlatency has been described with respect to a system having secure andnon-secure domains and a monitor mode, it is clearly applicable to anysystem which has functions that should not be resumed for a particularreason. Generally such functions operate by disabling interrupts whichincrease interrupt latency. Amending the functions to be restartable andcontrolling the processor to restart them following an interrupt allowsthe interrupts to be enabled for at least a portion of the processing ofthe function and helps reduce interrupt latency. For example normalcontext switching of an operating system.

Access to Secure and Non-Secure Memory

[0273] As described with reference to FIG. 1, the data processingapparatus has memory, which includes, inter alia, the TCM 36, cache 38,ROM 44, memory of slave devices and external memory 56. As describedwith reference to FIG. 37 for example, memory is partitioned into secureand non-secure memory. It will be appreciated that there will nottypically be any physical distinction between the secure memory regionsand non-secure memory regions of the memory at the time of fabrication,but that these regions will instead be defmed by a secure operatingsystem of the data processing apparatus when operating in the securedomain. Hence, any physical part of the memory device may be allocatedas secure memory, and any physical part may be allocated as non-securememory.

[0274] As described with reference to FIGS. 2 to 5, the processingsystem has a secure domain and a non-secure domain. In the securedomain, a secure kernel program 80 is provided and which executes in asecure mode. A monitor program 72 is provided which straddles the secureand non-secure domains and which executes at least partly in a monitormode. In embodiments of the invention the monitor program executespartly in the monitor mode and partly in a secure mode. As shown in forexample FIG. 10, there are a plurality of secure modes including, interalia, a supervisor mode SVC.

[0275] The monitor program 72 is responsible for managing all changesbetween the secure and non-secure domains in either direction. Some ofits functions are described with reference to FIGS. 8 and 9 in thesection ‘Processor Modes’. The monitor program is responsive to a modeswitching request SMI issued in the non-secure mode to initiate a switchfrom the said non-secure mode to the said secure mode and to a modeswitching request SMI issued in the secure mode to initiate a switchfrom the said secure mode to the said non-secure mode. As described inthe section ‘Switching between worlds’, in the monitor mode, switchingtakes place switching at least some of the register from one of thesecure and non-secure domains to the other. That involves saving thestate of a register existing in one domain and writing a new state tothe register (or restoring a previously saved state in the register) inthe other domain. As also described herein access to some registers maybe disabled when performing such a switch. Preferably, in monitor modeall interrupts are disabled.

[0276] Because the monitor mode in which the monitor program executesstraddles the secure and non-secure domains it is important that themonitor program is provably secure: that is it implements only thosefunctions it is intended to implement. It is thus advantageous if themonitor program is a simple as possible. The secure modes allowprocesses to execute only in the secure domain. In this embodiment ofthe present invention, the privileged secure mode(s) and the monitormode allows access to the same secure and non-secure memory. By ensuringthat the privileged secure mode(s) ‘see’ the same secure and non-securememory, functions which could otherwise only be implemented in themonitor mode are transferred to the secure mode allowing simplificationof the monitor program. In addition, this allows a process operating ina privileged secure mode to switch directly to monitor mode and viceversa. A switch from a privileged secure mode to the monitor mode ispermitted and in the monitor mode a switch to the non-secure domain maybe made. Non-privileged secure modes must use an SMI to enter themonitor mode. The system enters the privileged secure mode following areset. Switches between the monitor mode and the privileged secure modeand back are made to facilitate state saving when moving betweendomains.

[0277] In other embodiments access to the S flag may be allowed fromwithin secure privileged modes as well as from within the monitor mode.If secure privileged modes are allowed to switch the processor intomonitor mode whilst maintaining control of the program flow, then suchsecure privileged modes already effectively have the ability to changethe S flag (bit). Thus, the additional complexity of providing that theS flag can only be changed within the monitor mode is not justified. TheS flag can instead be stored in the same way as other configurationflags which may be changed by one or more secure privileged modes. Suchembodiments where the S flag may be changed within one of more secureprivileged modes are included within the current techniques.

[0278] Returning to the previously discussed example embodiment, theapparatus has a processor core 10 which defines the modes and definesthe privilege levels of the modes; i.e. the set of functions which anymode allows. Thus the processor core 10 is arranged in known manner toallow the secure modes and the monitor mode access to secure andnon-secure memory and the secure modes access to all memory to which themonitor mode allows access and to allow a process operating in anyprivileged secure mode to switch directly to monitor mode and viceversa. The processor core 10 is preferably arranged to allow thefollowing.

[0279] In one example of the apparatus, the memory is partitioned intosecure memory and non-secure memory, and both secure and non-securememory is accessible only in the monitor and secure modes. Preferably,the non-secure memory is accessible in monitor mode, a secure mode and anon-secure mode.

[0280] In another example of the apparatus, in the monitor mode and oneor more of the secure modes, access to the non-secure memory is deniedto the secure mode; and in non-secure mode access to the non-securememory is denied to the secure and monitor modes. Thus secure memory isaccessed only in monitor and secure modes and non-secure memory isaccessed only by non-secure modes increasing security.

[0281] In examples of the apparatus, resetting or booting of theapparatus may be performed in the monitor mode which may be regarded asa mode which is more privileged than a secure mode. privileged mode.However, in many examples of the apparatus are arranged to provideresetting or booting in a secure mode which is possible because of thedirect switching allowed between the secure mode and the monitor mode.

[0282] As described with reference to FIG. 2, in the secure domain, andin a secure mode, a secure kernel 80 (or operating system) functions,and one or more secure application programs 82, 84 may be run under thesecure kernel 80. The secure kernel and/or the secure applicationprogram or any other program code running in a secure mode is allowedaccess to both secure and non-secure memory.

[0283] Whilst examples of this invention have been described withreference to apparatus having a processor, the invention may beimplemented by a computer program which when run on a suitable processorconfigure s the processor to operate as described in this section.

[0284] A description of an alternative embodiment(s) of the presenttechnique considered from a programmer's model view is given below inrelation to FIGS. 21 to 23 as follows:

[0285] In the following description, we will use the following termsthat must be understood in the context of an ARM processor as designedby ARM Limited, of Cambridge, England.

[0286] S bit: Secure state bit, contained in a dedicated CP15 register.

[0287] ‘Secure/Non-Secure state’. This state is defined by the S bitvalue. It indicates whether the core may access the Secure world (whenit is in Secure state, i.e. S=1) or is restricted to the Non-secureworld only (S=0). Note that the Monitor mode (see further) overrides theS bit status.

[0288] ‘Non-Secure World’ groups all hardware/software accessible tonon-secure applications that do not require security.

[0289] ‘Secure World’ groups all hardware/software (core, memory . . . )that is only accessible when we execute secure code.

[0290] Monitor mode: new mode that is responsible for switching the corebetween the Secure and Non-secure state.

[0291] As a brief summary

[0292] The core can always access the Non-secure world.

[0293] The core can access the Secure world only when it is in Securestate or Monitor mode.

[0294] SMI: Software Monitor Interrupt: New instruction that will makethe core enter the Monitor mode through a dedicated SMI exceptionvector. ‘Thread ID’: is the identifier associated to each thread(controlled by an OS). For some types of OS where the OS runs innon-secure world, each time a secure function is called, it will benecessary to pass as a parameter the current thread ID to link thesecure function to its calling non-secure application. The secure worldcan thus support multi-threads.

[0295] Secure Interrupt defines an interrupt generated by a Secureperipheral.

Programmer's Model Carbon Core Overview

[0296] The concept of the Carbon architecture, which is the term usedherein for processors using the present techniques, consists inseparating two worlds, one secure and one non-secure. The secure worldmust not leak any data to non-secure world.

[0297] In the proposed solution, the secure and non-secure states willshare the same (existing) register bank. As a consequence, all currentmodes present in ARM cores (Abort, Undef, Irq, User, . . . ) will existin each state.

[0298] The core will know it operates in secure or non-secure statethanks to a new state bit, the S (secure) bit, instantiated in adedicated CP15 register.

[0299] Controlling which instruction or event is allowed to modify the Sbit, i.e. to change from one state to the other, is a key feature of thesecurity of the system. The current solution proposes to add a new mode,the Monitor mode, that will “supervise” switching between the twostates. The Monitor mode, by writing to the appropriate CP15 register,would be the only one allowed to alter the S bit.

[0300] Finally, we propose to add some flexibility to the exceptionhandling. All exceptions, apart from the reset, would be handled eitherin the state where they happened, or would be directed to the Monitormode. This would be left configurable thanks to a dedicated CP15register.

[0301] The details of this solution are discussed in the followingparagraphs.

Processor State and Modes Carbon New Features Secure or Non-Secure State(S Bit)

[0302] One major feature of the Carbon core is the presence of the Sbit, which indicates whether the core is in a Secure (S=1) or Non-secure(S=0) state. When in Secure state, the core would be able to access anydata in the Secure or Non-secure worlds. When in Non-Secure state, thecore would be restricted to the Non-secure world only.

[0303] The only exception to this rule concerns the Monitor mode, whichoverrides the S bit information. Even when S=0, the core will performSecure privileged accesses when it is in Monitor mode. See nextparagraph, Monitor mode, for further information

[0304] The S bit can only be read and written in Monitor mode. Whateverthe S bit value, if any other mode tries to access it, this will beeither ignored or result in an Undefined exception.

[0305] All exceptions, apart from Reset, have no effect on the Securestate bit. On Reset, the S bit will be set, and the core will start inSupervisor mode. Refer to the boot section for detailed information.

[0306] Secure/Nonsecure states are separate and operate independently ofthe ARM/Thumb/Java states.

Monitor Mode

[0307] One other important feature of the Carbon system is the creationof a new mode, the Monitor mode. This will be used to control the coreswitching between the Secure and Non-secure states. It will always beconsidered as a secure mode, i.e. whatever the value of the S bit, thecore will always perform Secure Privileged accesses to the externalworld when it is in Monitor mode.

[0308] Any Secure privileged mode (i.e. privileged modes when S=1) wouldbe able to switch to Monitor mode by simply writing the CPSR mode bits(MSR, MOVS, or equivalent instruction). However, this would be forbiddenin any Non-secure mode or Secure user mode. If this ever happens, theinstruction would be ignored or cause an exception.

[0309] There may be a need for a dedicated CPSR violation exception.This exception would be raised by any attempt to switch to Monitor modeby directly writing the CPSR from any Non-secure mode or Secure usermode.

[0310] All exceptions except Reset are in effect disabled when Monitormode is active:

[0311] all interrupts are masked;

[0312] all memory exceptions are either ignored or cause a fatalexception.

[0313] undefined/SWI/SMI are ignored or cause a fatal exception.

[0314] When entering Monitor mode, the interrupts are automaticallydisabled and the system monitor should be written such that none of theother types of exception can happen while the system monitor is running.

[0315] Monitor mode needs to have some private registers. This solutionproposes that we only duplicate the minimal set of registers, i.e R13(sp_mon), R14 (lr_mon) and SPSR (spsr_mon).

[0316] In Monitor mode, the MMU will be disabled (flat address map) aswell as the MPU or partition checker (the Monitor mode will alwaysperform secure privileged external accesses). However, speciallyprogrammed MPU region attributes (cacheability, . . . ) would still beactive. As an alternative the Monitor mode may use whatever mapping isused by the secure domain.

New Instruction

[0317] This proposal requires adding one new instruction to the existingARM instruction set.

[0318] The SMI (Software Monitor Interrupt) instruction would be used toenter the Monitor mode, branching at a fixed SMI exception vector. Thisinstruction would be mainly used to indicate to the Monitor to swapbetween the Non-secure and Secure State.

[0319] As an alternative (or in addition) it would be possible to add anew instruction to allow the Monitor mode to save/restore the state ofany other mode onto/from the Monitor stack to improve context switchingperformance.

Processor Modes

[0320] As discussed in the previous paragraph, only one new mode isadded in the core, the Monitor mode. All existing modes remainavailable, and will exist both in the secure and non-secure states.

[0321] In fact, Carbon users will see the structure illustrated in FIG.21.

Processor Registers

[0322] This embodiment proposes that the secure and the non-secureworlds share the same register bank. This implies that, when switchingfrom one world to the other through the Monitor mode, the system monitorwill need to save the first world context, and create (or restore) acontext in the second world.

[0323] Passing parameters becomes an easy task: any data contained in aregister in the first world will be available in the same register inthe second world once the system monitor has switched the S bit.

[0324] However, apart from a limited number of registers dedicated topassing parameters, which will need to be strictly controlled, all otherregisters will need to be flushed when passing from Secure to Non-securestate in order to avoid any leak of Secure data. This will need to beensured by the Monitor kernel.

[0325] The possibility of implementing a hardware mechanism or a newinstruction to directly flush the registers when switching from Secureto Non-secure state is also a possibility.

[0326] Another solution proposed involves duplicating all (or most of)the existing register bank, thus having two physically separatedregister banks between the Secure and Non-secure state. This solutionhas the main advantage of clearly separating the secure and non-securedata contained in the registers. It also allows fast context switchingbetween the secure and non-secure states. However, the drawback is thatpassing parameters through registers becomes difficult, unless we createsome dedicated instructions to allow the secure world access thenon-secure registers

[0327]FIG. 22 illustrates the available registers depending on theprocessor mode. Note that the processor state has no impact on thistopic.

Exceptions Secure Interrupts Current Solution

[0328] It is currently proposed to keep the same interrupt pins as inthe current cores, i.e. IRQ and FIQ. In association with the ExceptionTrap Mask register (defined later in the document), there should besufficient flexibility for any system to implement and handle differentkind of interrupts.

VIC Enhancement

[0329] We could enhance the VIC (Vectored Interrupt Controller) in thefollowing way: the VIC may contain one Secure information bit associatedto each vectored address. This bit would be programmable by the Monitoror Secure privileged modes only. It would indicate whether theconsidered interrupt should be treated as Secure, and thus should behandled on the Secure side.

[0330] We would also add two new Vector Address registers, one for allSecure Interrupts happening in Non-Secure state, the other one for allNon-Secure interrupts happening in Secure state.

[0331] The S bit information contained in CP15 would be also availableto the VIC as a new VIC input.

[0332] The following table summarizes the different possible scenarios,depending on the status of the incoming interrupt (Secure or Non-secure,indicated by the S bit associated to each interrupt line) and the stateof the core (S bit in CP15=S input signal on the VIC). Core in securestate Core in Non-secure state (CP15 − S=1) (CP15 − S=0) Secure No needto switch between worlds. The VIC has no Vector associated Interrupt TheVIC directly presents to the core the to this interrupt in theNon-secure Secure address associated to the interrupt line. domain. Itthus presents to the core The core simply has to branch at this addressthe address contained in the Vector where it should find the associatedISR. address register dedicated to all Secure interrupts occurring inNon- secure world. The core, still in Non- secure world, then branchesto this address, where it should find an SMI instruction to switch toSecure world. Once in Secure world, it would be able to have access tothe correct ISR. Non- The VIC has no Vector associated to this No needto switch between worlds. Secure interrupt in the Secure domain. It thusThe VIC directly presents to the Interrupt presents to the core theaddress contained in core the Non-secure address the Vector addressregister dedicated to all associated to the interrupt line. TheNon-secure interrupts occurring in Secure core simply has to branch atthis world. The core, still in Secure world, then address where itshould find the branches to this address, where it should findassociated Non-secure ISR. an SMI instruction to switch to Non-secureworld. Once in Non-secure world, it would be able to have access to thecorrect ISR.

Exception Handling Configurability

[0333] In order to improve Carbon flexibility, a new register, theException Trap Mask, would be added in CP15. This register would containthe following bits: Bit 0: Undef exception (Non-secure state) Bit 1: SWIexception (Non-secure state) Bit 2: Prefetch abort exception (Non-securestate) Bit 3: Data abort exception (Non-secure state) Bit 4: IRQexception (Non-secure state) Bit 5: FIQ exception (Non-secure state) Bit6: SMI exception (both Non-secure/Secure states) Bit 16: Undef exception(Secure state) Bit 17: SWI exception (Secure state) Bit 18: Prefetchabort exception (Secure state) Bit 19: Data abort exception (Securestate) Bit 20: IRQ exception (Secure state) Bit 21: FIQ exception(Secure state)

[0334] The Reset exception does not have any corresponding bit in thisregister. Reset will always cause the core to enter the Securesupervisor mode through its dedicated vector.

[0335] If the bit is set, the corresponding exception makes the coreenter the Monitor mode. Otherwise, the exception will be handled in itscorresponding handler in the world where it occurred.

[0336] This register would only be visible in Monitor mode. Anyinstruction trying to access it in any other mode would be ignored.

[0337] This register should be initialized to a system-specific value,depending upon whether the system supports a monitor or not. Thisfunctionality could be controlled by a VIC.

Exception Vectors Tables

[0338] As there will be separate Secure and Non-secure worlds, we willalso need separate Secure and Non-secure exception vectors tables.

[0339] Moreover, as the Monitor can also trap some exceptions, we mayalso need a third exception vectors table dedicated to the Monitor.

[0340] The following table summarizes those three different exceptionvectors tables:

[0341] In non-secure memory: Address Exception Mode Automaticallyaccessed when 0x00 — — 0x04 Undef Undef Undefined instruction executedwhen core is in Non-Secure state and Exception Trap Mask reg [Non-secureUndef]=0 0x08 SWI Supervisor SWI instruction executed when core is inNon-Secure state and Exception Trap Mask reg [Non-secure SWI]=0 0x0CPrefetch Abort Aborted instruction when core is in Abort Non-Securestate and Exception Trap Mask reg [Non-secure PAbort]=0 0x10 Data AbortAborted data when core is in Non- Abort Secure state and Exception TrapMask reg [Non-secure DAbort]=0 0x14 Reserved 0x18 IRG IRQ IRG pinasserted when core is in Non- Secure state and Exception Trap Mask reg[Non-secure IRQ]=0 0x1C FIQ FIQ FIQ pin asserted when core is in Non-Secure state and Exception Trap Mask reg [Non-secure FIQ]=0

[0342] In secure memory: Address Exception Mode Automatically accessedwhen 0x00 Reset* Supervisor Reset pin asserted 0x04 Undef UndefUndefined instruction executed when core is in Secure state andException Trap Mask reg [Secure Undef]=0 0x08 SWI Supervisor SWIinstruction executed when core is in Secure state and Exception TrapMask reg [Secure SWI]=0 0x0C Prefetch Abort Aborted instruction whencore Abort is in Secure state and Exception Trap Mask reg [SecurePAbort]=0 0x10 Data Abort Aborted data when core is in Secure Abortstate and Exception Trap Mask reg [Secure Dabort]=0 0x14 Reserved 0x18IRQ IRQ IRQ pin asserted when core is in Secure state and Exception TrapMask reg [Secure IRQ]=0 0x1C FIQ FIQ FIQ pin asserted when core is inSecure state and Exception Trap Mask reg [Secure FIQ]=0

[0343] In Monitor memory (flat mapping): Address Exception ModeAutomatically accessed when 0x00 — — — 0x04 Undef Monitor Undefinedinstruction executed when Core is in Secure state and Exception TrapMask reg [Secure Undef]=1 Core is in Non-secure state and Exception TrapMask reg [Non-secure Undef]=1 0x08 SWI Monitor SWI instruction executedwhen Core is in Secure state and Exception Trap Mask reg [Secure SWI]=1Core is in Non-secure state and Exception Trap Mask reg [Non-secureSWI]=1 0x0C Prefetch Monitor Aborted instruction when Abort Core is inSecure state and Exception Trap Mask reg [Secure IAbort]=1 Core is inNon-secure state and Exception Trap Mask reg [Non-secure Iabort]=1 0x10Data Monitor Aborted data when Abort Core is in Secure state andException Trap Mask reg [Secure PAbort]=1 Core is in Non-secure stateand Exception Trap Mask reg [Non-secure Pabort]=1 0x14 SMI Monitor 0x18IRQ Monitor IRQ pin asserted when core is in Secure state and ExceptionTrap Mask reg [Secure IRQ]=1 core is in Non-secure state and ExceptionTrap Mask reg [Non-secure IRQ]=1 0x1C FIQ Monitor FIQ pin asserted whencore is in Secure state and Exception Trap Mask reg [Secure FIQ]=1 coreis in Non-secure state and Exception Trap Mask reg [Non-secure FIQ]=1

[0344] In Monitor mode, the exceptions vectors may be duplicated, sothat each exception will have two different associated vector:

[0345] One for the exception arising in Non-secure state

[0346] One for the exception arising in Secure state

[0347] This may be useful to reduce the exception latency, because themonitor kernel does not have any more the need to detect the originatingstate where the exception occurred.

[0348] Note that this feature may be limited to a few exceptions, theSMI being one of the most suitable candidates to improve the switchingbetween the Secure and Non-secure states.

Switching Between Worlds

[0349] When switching between states, the Monitor mode must save thecontext of the first state on its Monitor stack, and restore the secondstate context from the Monitor stack.

[0350] The Monitor mode thus needs to have access to any register of anyother modes, including the private registers (r14, SPSR, . . . ).

[0351] To handle this, the proposed solution consists in giving anyprivilege mode in Secure state the rights to directly switch to Monitormode by simply writing the CPSR.

[0352] With such a system, switching between worlds is performed asfollows:

[0353] enter Monitor mode

[0354] set the S bit

[0355] switch to supervisor mode—save the supervisor registers on theMONITOR stack (of course the supervisor mode will need to have access tothe Monitor stack pointer, but this can be easily done, for example byusing a common register (R0 to R8))

[0356] switch to System mode—save the registers (=same as the user mode)on the Monitor stack

[0357] IRQ registers on the Monitor stack etc . . . for all modes

[0358] Once all private registers of all modes are saved, revert toMonitor mode with a simple MSR instruction (=simply write Monitor valuein the CPSR mode field)

[0359] The other solutions have also been considered:

[0360] Add a new instruction that would allow the Monitor to save othermodes' private registers on its own stack.

[0361] Implement the Monitor as a new “state”, i.e. being able to be inMonitor state (to have the appropriate access rights) and in IRQ (or anyother) mode, to see the IRQ (or any other) private registers.

[0362] Basic Scenario (See FIG. 23)

[0363] 1. Thread 1 is running in non-secure world (S bit=0) This threadneeds to perform a secure function=>SMI instruction.

[0364] 2. The SMI instruction makes the core enter the Monitor modethrough a non-secure SMI vector.

[0365] LR_mon and SPSR_mon are used to save the PC and CPSR of the nonsecure mode.

[0366] At this stage the S bit remains unchanged, although the system isnow in a secure state.

[0367] The monitor kernel saves the non-secure context on the monitor.

[0368] It also pushes LR_mon and SPSR_mon.

[0369] The monitor kernel then changes the “S” bit by writing into theCP15 register.

[0370] In this embodiment the monitor kernel keeps track that a “securethread 1” will be started in the secure world (e.g. by updating a ThreadID table).

[0371] Finally, it exits the monitor mode and switches to securesupervisor mode (MOVS instruction after having updated LR_mon andSPSR_mon?).

[0372] 3. The secure kernel dispatches the application to the rightsecure memory location, then switches to user mode (e.g. using a MOVS).

[0373] 4. The secure function in executed in secure user mode. Oncefinished, it calls an “exit” function by performing an appropriate SWI.

[0374] 5. The SWI instruction makes the core enter the secure svc modethrough a dedicated SWI vector, that in turn performs the “exit”function. This “exit” function ends with an “SMI” to switch back tomonitor mode.

[0375] 6. The SMI instruction makes the core enter the monitor modethrough a dedicated secure SMI vector.

[0376] LR_mon and SPSR_mon are used to save the PC and CPSR of theSecure svc mode.

[0377] The S bit remains unchanged (i.e. Secure State).

[0378] The monitor kernel registers the fact that secure thread 1 isfinished (removes the secure thread 1 ID from the thread ID table?).

[0379] It then changes the “S” bit by writing into the CP15 register,returning to non-secure state.

[0380] The monitor kernel restores the non-secure context from themonitor stack.

[0381] It also load the LR_mon and CPSR mon previously saved in step 2.

[0382] Finally, it exits monitor mode with a SUBS, that will make thecore return in non-secure user mode, on the instruction

[0383] 7. Thread 1 can resume normally.

[0384] Referring to FIG. 6, all of the registers are shared between thesecure and non-secure domains In monitor mode, switching takes placeswitching the registers from one of the secure and non-secure domains tothe other. That involves saving the state of a register existing in onedomain and writing a new state to the register (or restoring apreviously saved state in the register) in the other domain as is alsodescribed in the section ‘Switching between Worlds’ above.

[0385] It is desirable to reduce the time taken to perform this switch.To reduce the time taken to perform the switch, the shared registers aredisabled when switching between the secure and non-secure domainsretaining unchanged the data values stored therein. For example,consider a switch from the non-secure domain to the secure domain.Assume that for example the FIQ registers shown in FIG. 6 are not neededin the secure world. Thus, those registers are disabled and there is noneed to switch them to the secure domain and there is no need to savethe contents of those registers.

[0386] Disabling of the registers may be achieved in several ways. Oneway is to lock out the mode which uses those registers. That is done bywriting a control bit in a CP15 register indicating the disabling ofthat mode.

[0387] Alternatively, access to the registers may be disabled on aninstruction by instruction basis again by writing control bits in a CP15register. The bits written in the CP15 register relate explicitly to theregister, not the mode, so the mode is not disabled but access to theregister in the mode is disabled.

[0388] The FIQ registers store data associated with a fast interrupt. Ifthe FIQ register(s) are disabled and a fast interrupt occurs, theprocessor signals an exception in the monitor. In response to anexception, the monitor mode is operable to save any data valuesassociated with one domain and stored in the said disabled register andto load into that register new data values associated with the otherdomain and then re-enable the FIQ mode registers.

[0389] The processor may be arranged so that when in the monitor modeall banked registers are disabled when the processor switches domains.Alternatively, the disabling of the registers may be selective in thatsome predetermined ones of the shared registers are disabled whenswitching domains and others may be disabled at the choice of theprogrammer.

[0390] The processor may be arranged so that when switching domains inthe monitor mode, one or more of the shared registers are disabled, andone or more others of the shared registers have their data saved whenexisting one domain, and have new data loaded in the other domain. Thenew data may be null data.

[0391]FIG. 24 schematically illustrates the concept of adding a secureprocessing option to a traditional ARM core. The diagram schematicallyshows how a processor that contains a secure processing option can beformed by adding a secure processing option to an existing core. If thesystem is to be backwards compatible with an existing legacy operatingsystem, it is intuitive to think of the legacy system operating in thetraditional non-secure part of the processor. However, as is shownschematically in the lower part of the Figure and is detailed furtherbelow, it is in fact in the secure portion of the system that a legacysystem operates.

[0392]FIG. 25 shows a processor having a secure and non-secure domainand illustrating reset and is similar to FIG. 2. FIG. 2 illustrates aprocessor that is adapted to run a security sensitive type of operationwith a secure OS system controlling processing in the secure domain anda non-secure OS system controlling processing in the non-secure domain.The processor is however also backwards compatible with a traditionallegacy operating system and thus, the processor may operate in asecurity insensitive way using a legacy operating system.

[0393] As is shown in FIG. 25, the reset is in the secure domain, andwhatever the type of operation reset occurs here with the S-bit orsecurity status flag set. In the case of a security insensitive type ofoperation, reset occurs in the secure domain and processing thencontinues within the secure domain. The legacy operating systemcontrolling processing is however unaware of the security aspects of thesystem.

[0394] As is shown in FIG. 25 reset is performed to set the address atwhich to start processing within the secure supervisor mode whetherprocessing is to be secure sensitive or is in fact secure insensitive.Once reset has been performed the additional tasks present in a boot orreboot mechanism are then performed. The boot mechanism is describedbelow.

Boot Mechanism

[0395] The boot mechanism must respect the following features:

[0396] Keep compatibility with legacy OSes.

[0397] Boot in most privileged mode to ensure the security of thesystem.

[0398] As a consequence, Carbon cores will boot in Secure Supervisormode.

[0399] The different systems will then be:

[0400] For systems wanting to run legacy OSes, the S bit is not takeninto account and the core will just see it boots in Supervisor mode.

[0401] For systems wanting to use the Carbon features, the core boots inSecure privileged mode which should be able to configure all secureprotections in the system (potentially after swapping to Monitor mode)

[0402] With respect to the details of the boot mechanism given above theprocessor of embodiments of the present invention resets the processorto start processing in the secure supervisor mode in all cases. In thecase of a security insensitive type of operation the operating system isin effect operating in the secure domain although security is not herean issue, because the S bit is set (although the operating system isunaware of this). This has the advantage that parts of the memory thatare inaccessible from the non-secure domain are accessible in thissituation.

[0403] Booting in secure supervisor mode in all cases is alsoadvantageous in security sensitive systems as it helps ensure thesecurity of the system. In secure sensitive systems, the addressprovided at boot points to where the boot program is stored in securesupervisor mode and thus, enables the system to be configured as asecure system and to switch to monitor mode. Switching from securesupervisor mode to monitor mode is allowed in general and enables thesecure system at an appropriate time to start processing in monitor modeto initialise monitor mode configuration.

[0404]FIG. 26 illustrates at step 1 a non-secure thread NSA beingexecuted by a non-secure operating system. At step 2 the non-securethread NSA makes a call to the secure domain via the monitor moderunning a monitor mode program at step 3. The monitor mode programchanges the S-bit to switch domain and performs any necessary contextsaving and context restoring prior to moving to the secure operatingsystem at step 5. The corresponding secure thread SA is then executedbefore it is subject to an interrupt irq at step 6. The interrupthandling hardware triggers a return to the monitor mode at step 7 whereit is determined as to whether the interrupt will be handled by thesecure operating system or the non-secure operating system. In thiscase, the interrupt is handled by the non-secure operating systemstarting at step 9. When this interrupt has been handled by thenon-secure operating system, the non-secure thread NSA is resumed as thecurrent task in the non-secure operating system prior to a normal threadswitching operation at step 11. This thread switching may be the resultof a timing event or the like. A different thread NSB is executed in thenon-secure domain by the non-secure operating system at step 12 and thisthen makes a call to the secure domain via the monitor domain/program atstep 14. The monitor program at step 7 has stored a flag, used someother mechanism, to indicate that the secure operating system was lastsuspended as a result of an interrupt rather than having been leftbecause a secure thread had finished execution or due to a normalrequest to leave. Accordingly, since the secure operating system wassuspended by an interrupt, the monitor program at step 15 re-enters thesecure operating system using a software faked interrupt which specifiesa return thread ID (e.g. an identifier of the thread to be started bythe secure operating system as requested by the non-secure thread NSB,as well as other parameter data). These parameters of the software fakedinterrupt may be passed as register values.

[0405] The software faked interrupt triggers a return interrupt handlerroutine of the secure operating system at step 15. This return interrupthandler routine examines the return thread ID of the software fakedinterrupt to determine whether or not this matches the thread ID of thesecure thread SA which was interrupted the last time the secureoperating system was executed prior to its suspension. In this case,there is not a match and accordingly at step 16 the secure operatingsystem is triggered to perform a thread switch to the return thread asspecified by the non-secure thread NSB after the context of the securethread SA has been saved. The secure thread SA can then later berestarted from the point at which it was interrupted as required.

[0406]FIG. 27 schematically illustrates another example of the type ofbehaviour illustrated in FIG. 26. In this example whilst processingproceeds under control of the non-secure operating system to handle theirq, there is no non-secure thread switch and accordingly when thesoftware faked interrupt is received by the return interrupt handler ofthe secure operating system it determines that no thread switch isrequired and simply resumes the secure thread SA at step 15.

[0407]FIG. 28 is a flow diagram schematically illustrating theprocessing performed by the return thread handler. At step 4002 thereturn thread handler is started. At step 4004 the return threadidentifier from the software faked interrupt is examined and comparedwith the currently executing secure thread when the secure operatingsystem was suspended. If these match, then processing proceeds to step4006 at which the secure thread is resumed. If the comparison at step4004 is not matched, then processing proceeds to step 4008 at which thecontext of the old secure thread is saved, (for subsequent resumption)prior to a switch being made to the new secure thread at step 4010. Thenew thread might already be under way and so step 4010 is a resumption.

[0408]FIG. 29 schematically illustrates processing whereby a slavesecure operating system may follow task switches performed by a masternon-secure operating system. The master non-secure operating system maybe a legacy operating system with no mechanisms for communicating andco-ordinating its activities to other operating systems and accordinglyoperate only as a master. As an initial entry point in FIG. 29 thenon-secure operating system is executing a non-secure thread NSA. Thisnon-secure thread NSA makes a call to a secure thread which is to beexecuted by the secure operating system using a software interrupt, anSMI call. The SMI call goes to a monitor program executing in a monitormode at step 2 whereupon the monitor program performs any necessarycontext saving and switching before passing the call onto the secureoperating system at step 4. The secure operating system then starts thecorresponding secure thread SA. This secure thread may return controlvia the monitor mode to the non-secure operating system, such as as aresult of a timer event or the like. When the non-secure thread NSAagain passes control to the secure operating system at step 9 it does soby reissuing the original software interrupt. The software interruptincludes the non-secure thread ID identifying NSA, the secure thread IDof the target secure thread to be activated, i.e. the thread IDidentifying secure thread SA as well as other parameters.

[0409] When the call generated at step 9 is passed on by the monitorprogram and received at step 12 in the secure domain by the secureoperating system, the non-secure thread ID can be examined to determinewhether or not there has been a context switch by the non-secureoperating system. The secure thread ID of the target thread may also beexamined to see that the correct thread under the secure operatingsystem is restarted or started as a new thread. In the example of FIG.29, no thread switch is required in the secure domain by the secureoperating system.

[0410]FIG. 30 is similar to FIG. 29 except that a switch in threadoccurs at step 9 in the non-secure domain under control of thenon-secure operating system. Accordingly, it is a different non-securethread NSB which makes the software interrupt call across to the secureoperating system at step 11. At step 14, the secure operating systemrecognises the different thread ID of the non-secure thread NSB andaccordingly performs a task switch involving saving the context of thesecure thread SA and starting the secure thread SB.

[0411]FIG. 31 is a flow diagram schematically illustrating processingperformed by the secure operating system when receiving a softwareinterrupt as a call to start a thread or resume a thread of the secureoperating system. At step 4012 the call is received. At step 4014 theparameters of the call are examined to determine if they match thecurrently active secure thread upon the secure operating system. If amatch occurs, then this secure thread is restarted at step 4016. If amatch does not occur, then processing proceeds to step 4018 where adetermination is made as to whether or not the newly requested thread isavailable. The newly requested thread might be unavailable due to areason such as it being or requiring an exclusive resource which isalready in use by some other thread executing on a secure operatingsystem. In such a case, the call is rejected at step 4020 with anappropriate message being returned to the non-secure operating system.If the determination at step 4018 is that the new thread is available,then processing proceeds to step 4022 at which the context of the oldsecure thread is saved for possible later resumption. At step 4024 aswitch is made to the new secure thread as specified in the softwareinterrupt call made to the secure operating system.

[0412]FIG. 32 schematically illustrates operation whereby a priorityinversion may occur when handling interrupts within a system havingmultiple operating systems with different interrupts being handled bydifferent operating systems.

[0413] Processing starts with the secure operating system executing asecure thread SA. This is then interrupted by a first interrupt Intl.This triggers the monitor program within the monitor mode to determinewhether or not the interrupt is to be handled in the secure domain orthe non-secure domain. In this case, the interrupt is to be handled inthe secure domain and processing is returned to the secure operatingsystem and the interrupt handling routine for interrupt Intl is started.Partway through execution of the interrupt handling routine for Int1, afurther interrupt Int2 is received which has a higher priority. Thus,the interrupt handler for Int 1 is stopped and the monitor program inthe monitor mode used to determine where the interrupt Int2 is to behandled. In this case the interrupt Int2 is to be handled by thenon-secure operating system and accordingly control is passed to thenon-secure operating system and the interrupt handler for Int2 started.When this interrupt handler for the interrupt Int2 has completed, thenon-secure operating system has no information indicating that there isa pending interrupt Int1 for which servicing has been suspended in thesecure domain. Accordingly, the non-secure operating system may performsome further processing, such as a task switch or the starting of adifferent non-secure thread NSB, whilst the original interrupt Int1remains unserviced.

[0414]FIG. 33 illustrates a technique whereby the problems associatedwith the operation of FIG. 32 may be avoided. When the interrupt Int1occurs, the monitor program passes this to the non-secure domain where astub interrupt handler is started. This stub interrupt handler isrelatively small and quickly returns processing to the secure domain viathe monitor mode and triggers an interrupt handler for the interruptInt1 within the secure domain. The interrupt Int1 is primarily processedwithin the secure domain and the starting of the stub interrupt handlerin the non-secure domain can be regarded as a type of placeholder toindicate to the non-secure domain that the interrupt is pending in thesecure domain.

[0415] The interrupt handler in the secure domain for the interrupt Int1is again subject to a high priority Int2. This triggers execution of theinterrupt handler for the interrupt Int2 in the non-secure domain asbefore. However, in this case, when that interrupt handler for Int2 hasfinished, the non-secure operating system has data indicating that thestub interrupt handler for interrupt Int1 is still outstanding andaccordingly will resume this stub interrupt handler. This stub interrupthandler will appear as if it were suspended at the point at which itmade its call back to the secure domain and accordingly this call willbe re-executed and thus the switch made to the secure domain. Once backin the secure domain, the secure domain can itself re-start theinterrupt handler for the interrupt Int1 at the point at which it wassuspended. When the interrupt handler for the interrupt Int1 hascompleted within the secure domain, a call is made back to thenon-secure domain to close down the stub interrupt handler in thenon-secure domain before the originally executing secure thread SA isresumed.

[0416]FIG. 34 schematically illustrates different types of interruptwith their associated priorities and how they may be handled. Highpriority interrupts may be handled using purely secure domain interrupthandlers providing there is no higher priority interrupt that is handledby the non-secure domain. Once there is an interrupt having a priorityhigher than subsequent interrupts and which is handled in the non-securedomain, then all lower interrupts must either be handled purely in thenon-secure domain or utilise the stub interrupt handler techniqueillustrated in FIG. 33 whereby the non-secure domain can keep track ofthese interrupts even though their main handling is occurring within thesecure domain.

[0417] As mentioned earlier, the monitor mode is used to performswitching between the secure domain and the non-secure domain. Inembodiments where registers are shared between the two differentdomains, this involves saving the state within those registers intomemory, and then loading the new state for the destination domain frommemory into those registers. For any registers which are not sharedbetween the two domains, the state need not be saved away, since thoseregisters will not be accessed by the other domain, and switchingbetween the states is implemented as a direct result of switchingbetween the secure and non-secure domains (i.e. the value of the S bitstored in one of the CP15 registers determines which of the non-sharedregisters are used)

[0418] Part of the state that needs to be switched whilst in the monitormode is the processor configuration data controlling access to memory bythe processor. Since within each domain there is a different view of thememory, for example the secure domain having access to secure memory forstoring secure data, this secure memory not being accessible from thenon-secure domain, it is clear that the processor configuration datawill need to be changed when switching between the domains.

[0419] As illustrated in FIG. 35, this processor configuration data isstored within the CP15 registers 34, and in one embodiment theseregisters are shared between the domains. Hence, when the monitor modeis switched between the secure domain and the non-secure domain, theprocessor configuration data currently in the CP15 registers 34 needs tobe switched out of the CP15 registers into memory, and processorconfiguration data relating to the destination domain needs to be loadedinto the CP15 registers 34.

[0420] Since the processor configuration data in the CP15 registerstypically has an immediate effect on the access to memory within thesystem, then it is clear that these settings would become immediatelyeffective as they are updated by the processor whilst operating in themonitor mode. However, this is undesirable since it is desirable for themonitor mode to have a static set of processor configuration data thatcontrol access to memory whilst in monitor mode.

[0421] Accordingly, as shown in FIG. 35, in one embodiment of thepresent invention monitor mode specific processor configuration data2000 is provided, which can be used to override the processorconfiguration data in the CP15 registers 34 whilst the processor isoperating in the monitor mode. This is achieved in the embodimentillustrated in FIG. 35 through the provision of a multiplexer 2010 whichreceives at its inputs both the processor configuration data stored inthe CP15 registers and the monitor mode specific processor configurationdata 2000. Furthermore, the multiplexer 2010 receives a control signalover path 2015 indicating whether the processor is currently operatingin the monitor mode or not. If the processor is not operating in themonitor mode, then the processor configuration data in the CP15registers 34 is output to the system, but in the event that theprocessor is operating in the monitor mode, the multiplexer 2010 insteadoutputs the monitor mode specific processor configuration data 2000 toensure that a consistent set of processor configuration data is appliedwhile the processor is operating in the monitor mode.

[0422] The monitor mode specific processor configuration data can behard-coded within the system, thereby ensuring that it cannot bemanipulated. However, it is possible that the monitor mode specificprocessor configuration data 2000 could be made programmable withoutcompromising security, provided that that monitor mode specificprocessor configuration data could only be modified by the processorwhen operating in a secure privileged mode. This would allow someflexibility as to the setting of the monitor mode specific processorconfiguration data. If the monitor mode specific processor configurationdata is arranged to be programmable, that configuration data can bestored in any appropriate place within the system, for example within aseparate set of registers within the CP15 registers 34.

[0423] Typically, the monitor mode specific processor configuration datawill be set so as to provide a very secure environment for operation ofthe processor in the monitor mode. Hence, in the above-describedembodiment, the monitor mode specific processor configuration data mayspecify that the memory management unit 30 is disabled whilst theprocessor is operating in the monitor mode, thereby disabling anyvirtual to physical address translation that might otherwise be appliedby the memory management unit. In such a situation, the processor willalways be arranged to directly issue physical addresses when issuingmemory access requests, i.e. flat mapping will be employed. This ensuresthat the processor can reliably access memory whilst operating in themonitor mode, irrespective of whether any virtual to physical addressmappings have been tampered with.

[0424] The monitor mode specific processor configuration data would alsotypically specify that the processor is allowed to access the securedata whilst the processor is operating in the monitor mode. This ispreferably specified by memory permission data taking the form of adomain status bit, this domain status bit having the same value thatwould be specified for the corresponding domain status bit (“S” bit)within the secure processor configuration data. Hence, irrespective ofthe actual value of the domain status bit stored within the CP15registers, that value will get overridden by the domain status bitspecified by the monitor mode specific processor configuration data, toensure that the monitor mode has access to secure data.

[0425] The monitor mode specific processor configuration data may alsospecify other data used to control access to parts of the memory. Forexample, the monitor mode specific processor configuration data mayspecify that the cache 38 is not to be used to access data whilst theprocessor is operating in the monitor mode.

[0426] In the embodiment described above, it has been assumed that allof the CP15 registers containing processor configuration data are sharedbetween the domains. However, in an alternative embodiment, a number ofthe CP15 registers are “banked”, so that for example there are tworegisters for storing a particular item of processor configuration data,one register being accessible in the non-secure domain and containingthe value of that item of processor configuration data for thenon-secure domain, and the other register being accessible in the securedomain and containing the value of that item of processor configurationdata for the secure domain.

[0427] One CP15 register that will not be banked is the one containingthe “S” bit, but in principle any of the other CP15 registers may bebanked if desired. In such embodiments, the switching of the processorconfiguration data by the monitor mode involves switching out of anyshared CP15 registers into memory the processor configuration datacurrently in those shared registers, and loading into those shared CP15registers the processor configuration data relating to the destinationdomain. For any banked registers, the processor configuration data neednot be stored away to memory, and instead the switching will occurautomatically as a result of changing the S bit value stored in therelevant shared CP15 register.

[0428] As mentioned earlier, the monitor mode processor configurationdata will include a domain status bit which overrides that stored in therelevant CP15 register but has the same value as that used for thedomain status bit used in the secure domain (i.e. an S bit value of 1 inthe above described embodiments). When a number of the CP15 registersare banked, this means that at least part of the monitor mode specificprocessor configuration data 2000 in FIG. 35 can be derived from thesecure processor configuration data stored in banked registers sincethose registers contents are not written out to memory during theswitching process.

[0429] Hence, as an example, since the monitor mode specific processorconfiguration data will specify a domain status bit to override thatwhich is otherwise used when not in monitor mode, and in preferredembodiments this has the same value as that used in the secure domain,this means that the logic that selects which of the banked CP15registers are accessible will allow the secure banked CP15 registers tobe accessed. By allowing the monitor mode to use this secure processorconfiguration data as the relevant part of the monitor mode specificprocessor configuration data, a saving in resource can be realised sinceit is no longer necessary to provide a separate set of registers forthose items of monitor mode specific processor configuration data.

[0430]FIG. 36 is a flow diagram illustrating the steps performed toswitch the processor configuration data when a transition between onedomain and the other is required. As mentioned previously, an SMIinstruction is issued in order to instigate the transition betweendomains. Accordingly, at step 2020, the issuance of an SMI instructionis awaited. When an SMI instruction is received, the processor proceedsto step 2030, where the processor begins running the monitor program inmonitor mode, this causing the monitor mode specific processorconfiguration data to be used as a result of the control signal on path2015 into the multiplexer 2010 causing the multiplexer to switch to thatmonitor mode specific processor configuration data. As mentionedearlier, this can be a self-contained set of data, or certain parts ofit may be derived from the secure processor configuration data stored inbanked registers.

[0431] Thereafter, at step 2040, the current state is saved from thedomain issuing the SMI instruction into memory, this including savingfrom any shared CP15 registers the state of the processor configurationdata relevant to that domain. Typically, there will be a portion ofmemory set aside for the storage of such state. Then, at step 2050, thestate pointer is switched to point to the portion of memory thatcontains the corresponding state for the destination domain. Hence,typically, there will be two portions of memory allocated for storingstate information, one allocated for storing the state for thenon-secure domain, and one allocated for storing the state for thesecure domain.

[0432] Once the state pointer has been switched at step 2050, that statenow pointed to by the state pointer is loaded into the relevant sharedCP15 registers at step 2060, this including loading in the relevantprocessor configuration data for the destination domain. Thereafter, atstep 2070, the monitor program is exited, as is the monitor mode, andthe processor then switches to the required mode in the destinationdomain.

[0433]FIG. 37 illustrates in more detail the operation of the memorymanagement logic 30 of one embodiment of the present invention. Thememory management logic consists of a Memory Management Unit (MMU) 200and a Memory Protection Unit (MPU) 220. Any memory access request issuedby the core 10 that specifies a virtual address will be passed over path234 to the MMU 200, the MMU 200 being responsible for performingpredetermined access control functions, more particularly fordetermining the physical address corresponding to that virtual address,and for resolving access permission rights and determining regionattributes.

[0434] The memory system of the data processing apparatus consists ofsecure memory and non-secure memory, the secure memory being used tostore secure data that is intended only to be accessible by the core 10,or one or more other master devices, when that core or other device isoperating in a secure mode of operation, and is accordingly operating inthe secure domain.

[0435] In the embodiment of the present invention illustrated in FIG.37, the policing of attempts to access secure data in secure memory byapplications running on the core 10 in non-secure mode is performed bythe partition checker 222 within the MPU 220, the MPU 220 being managedby the secure operating system, also referred to herein as the securekernel.

[0436] In accordance with preferred embodiments of the present inventiona non-secure page table 58 is provided within non-secure memory, forexample within a non-secure memory portion of external memory 56, and isused to store for each of a number of non-secure memory regions definedwithin that page table a corresponding descriptor. The descriptorcontains information from which the MMU 200 can derive access controlinformation required to enable the MMU to perform the predeterminedaccess control functions, and accordingly in the embodiment describedwith reference to FIG. 37 will provide information about the virtual tophysical address mapping, the access permission rights, and any regionattributes.

[0437] Furthermore, in accordance with the preferred embodiments of thepresent invention, at least one secure page table 58 is provided withinsecure memory of the memory system, for example within a secure part ofexternal memory 56, which again for a number of memory regions definedwithin the table provides an associated descriptor. When the processoris operating in a non-secure mode, the non-secure page table will bereferenced in order to obtain relevant descriptors for use in managingmemory accesses, whilst when the processor is operating in secure mode,descriptors from the secure page table will be used.

[0438] The retrieval of descriptors from the relevant page table intothe MMU proceeds as follows. In the event that the memory access requestissued by the core 10 specifies a virtual address, a lookup is performedin the micro-TLB 206 which stores for one of a number of virtual addressportions the corresponding physical address portions obtained from therelevant page table. Hence, the micro-TLB 206 will compare a certainportion of the virtual address with the corresponding virtual addressportion stored within the micro-TLB to determine if there is a match.The portion compared will typically be some predetermined number of mostsignificant bits of the virtual address, the number of bits beingdependent on the granularity of the pages within the page table 58. Thelookup performed within the micro-TLB 206 will typically be relativelyquick, since the micro-TLB 206 will only include a relatively few numberof entries, for example eight entries

[0439] In the event that there is no match found within the micro-TLB206, then the memory access request is passed over path 242 to the mainTLB 208 which contains a number of descriptors obtained from the pagetables. As will be discussed in more detail later, descriptors from boththe non-secure page table and the secure page table can co-exist withinthe main TLB 208, and each entry within the main TLB has a correspondingflag (referred to herein as a domain flag) which is settable to indicatewhether the corresponding descriptor in that entry has been obtainedfrom a secure page table or a non-secure page table. In any embodimentswhere all secure modes of operation specify physical addresses directlywithin their memory access requests, it will be appreciated that therewill not be a need for such a flag within the main TLB, as the main TLBwill only store non-secure descriptors.

[0440] Within the main TLB 208, a similar lookup process is performed todetermine whether the relevant portion of the virtual address issuedwithin the memory access request corresponds with any of the virtualaddress portions associated with descriptors in the main TLB 208 thatare relevant to the particular mode of operation. Hence, if the core 10is operating in non-secure mode, only those descriptors within the mainTLB 208 which have been obtained from the non-secure page table will bechecked, whereas if the core 10 is operating in secure mode, only thedescriptors within the main TLB that have been obtained from the securepage table will be checked.

[0441] If there is a hit within the main TLB as a result of thatchecking process, then the access control information is extracted fromthe relevant descriptor and passed back over path 242. In particular,the virtual address portion and the corresponding physical addressportion of the descriptor will be routed over path 242 to the micro-TLB206, for storage in an entry of the micro-TLB, the access permissionrights will be loaded into the access permission logic 202, and theregion attributes will be loaded into the region attribute logic 204.The access permission logic 202 and region attribute logic 204 may beseparate to the micro-TLB, or may be incorporated within the micro-TLB.

[0442] At this point, the MMU 200 is then able to process the memoryaccess request since there will now be a hit within the micro-TLB 206.Accordingly, the micro-TLB 206 will generate the physical address, whichcan then be output over path 238 onto the system bus 40 for routing tothe relevant memory, this being either on-chip memory such as the TCM36, cache 38, etc, or one of the external memory units accessible viathe external bus interface 42. At the same time, the access permissionlogic 202 will determine whether the memory access is allowed, and willissue an abort signal back to the core 10 over path 230 if it determinesthat the core is not allowed to access the specified memory location inits current mode of operation. For example, certain portions of memory,whether in secure memory or non-secure memory, may be specified as onlybeing accessible by the core when that core is operating in supervisormode, and accordingly if the core 10 is seeking to access such a memorylocation when in, for example, user mode, the access permission logic202 will detect that the core 10 does not currently have the appropriateaccess rights, and will issue the abort signal over path 230. This willcause the memory access to be aborted. Finally, the region attributelogic 204 will determine the region attributes for the particular memoryaccess, such as whether the access is cacheable, bufferable, etc, andwill issue such signals over path 232, where they will then be used todetermine whether the data the subject of the memory access request canbe cached, for example within the cache 38, whether in the event of awrite access the write data can be buffered, etc.

[0443] In the event that there was no hit within the main TLB 208, thenthe translation table walk logic 210 is used to access the relevant pagetable 58 in order to retrieve the required descriptor over path 248, andthen pass that descriptor over path 246 to the main TLB 208 for storagetherein. The base address for both the non-secure page table and thesecure page table will be stored within registers of CP15 34, and thecurrent domain in which the processor core 10 is operating, i.e. securedomain or non-secure domain, will also be set within a register of CP15,that domain status register being set by the monitor mode when atransition occurs between the non-secure domain and the secure domain,or vice versa. The content of the domain status register will bereferred to herein as the domain bit. Accordingly, if a translationtable walk process needs to be performed, the translation table walklogic 210 will know in which domain the core 10 is executing, andaccordingly which base address to use to access the relevant table. Thevirtual address is then used as an offset to the base address in orderto access the appropriate entry within the appropriate page table inorder to obtain the required descriptor.

[0444] Once the descriptor has been retrieved by the translation tablewalk logic 210, and placed within the main TLB 208, a hit will then beobtained within the main TLB, and the earlier described process will beinvoked to retrieve the access control information, and store it withinthe micro-TLB 206, the access permission logic 202 and the regionattribute logic 204. The memory access can then be actioned by the MMU200.

[0445] As mentioned earlier, in preferred embodiments, the main TLB 208can store descriptors from both the secure page table and the non-securepage table, but the memory access requests are only processed by the MMU200 once the relevant information is stored within the micro-TLB 206. Inpreferred embodiments, the transfer of information between the main TLB208 and the micro-TLB 206 is monitored by the partition checker 222located within the MPU 220 to ensure that, in the event that the core 10is operating in a non-secure mode, no access control information istransferred into the micro-TLB 206 from descriptors in the main TLB 208if that would cause a physical address to be generated which is withinsecure memory.

[0446] The memory protection unit is managed by the secure operatingsystem, which is able to set within registers of the CP15 34partitioning information defining the partitions between the securememory and the non-secure memory. The partition checker 222 is then ableto reference that partitioning information in order to determine whetheraccess control information is being transferred to the micro-TLB 206which would allow access by the core 10 in a non-secure mode to securememory. More particularly, in preferred embodiments, when the core 10 isoperating in a non-secure mode of operation, as indicated by the domainbit set by the monitor mode within the CP15 domain status register, thepartition checker 222 is operable to monitor via path 244 any physicaladdress portion seeking to be retrieved into the micro-TLB 206 from themain TLB 208 and to determine whether the physical address that wouldthen be produced for the virtual address based on that physical addressportion would be within the secure memory. In such circumstances, thepartition checker 222 will issue an abort signal over path 230 to thecore 10 to prevent the memory access from taking place.

[0447] It will be appreciated that in addition the partition checker 222can be arranged to actually prevent that physical address portion frombeing stored in the micro-TLB 206 or alternatively the physical addressportion may still be stored within the micro-TLB 206, but part of theabort process would be to remove that incorrect physical address portionfrom the micro-TLB 206, for example by flushing the micro-TLB 206.

[0448] Whenever the core 10 changes via the monitor mode between anon-secure mode and a secure mode of operation, the monitor mode willchange the value of the domain bit within the CP15 domain statusregister to indicate the domain into which the processor's operation ischanging. As part of the transfer process between domains, the micro-TLB206 will be flushed and accordingly the first memory access following atransition between secure domain and non-secure domain will produce amiss in the micro-TLB 206, and require access information to beretrieved from main TLB 208, either directly, or via retrieval of therelevant descriptor from the relevant page table.

[0449] By the above approach, it will be appreciated that the partitionchecker 222 will ensure that when the core is operating in thenon-secure domain, an abort of a memory access will be generated if anattempt is made to retrieve into the micro-TLB 206 access controlinformation that would allow access to secure memory.

[0450] If in any modes of operation of the processor core 10, the memoryaccess request is arranged to specify directly a physical address, thenin that mode of operation the MMU 200 will be disabled, and the physicaladdress will pass over path 236 into the MPU 220. In a secure mode ofoperation, the access permission logic 224 and the region attributelogic 226 will perform the necessary access permission and regionattribute analysis based on the access permission rights and regionattributes identified for the corresponding regions within thepartitioning information registers within the CP15 34. If the securememory location seeking to be accessed is within a part of secure memoryonly accessible in a certain mode of operation, for example secureprivileged mode, then an access attempt by the core in a different modeof operation, for example a secure user mode, will cause the accesspermission logic 224 to generate an abort over path 230 to the core inthe same way that the access permission logic 202 of the MMU would haveproduced an abort in such circumstances. Similarly, the region attributelogic 226 will generate cacheable and bufferable signals in the same waythat the region attribute logic 204 of the MMU would have generated suchsignals for memory access requests specified with virtual addresses.Assuming the access is allowed, the access request will then proceedover path 240 onto the system bus 40, from where it is routed to theappropriate memory unit.

[0451] For a non-secure access where the access request specifies aphysical address, the access request will be routed via path 236 intothe partition checker 222, which will perform partition checking withreference to the partitioning information in the CP15 registers 34 inorder to determine whether the physical address specifies a locationwithin secure memory, in which event the abort signal will again begenerated over path 230.

[0452] The above described processing of the memory management logicwill now be described in more detail with reference to the flow diagramsof FIGS. 39 and 40. FIG. 39 illustrates the situation in which theprogram running on the core 10 generates a virtual address, as indicatedby step 300. The relevant domain bit within the CP15 domain statusregister 34 as set by the monitor mode will indicate whether the core iscurrently running in a secure domain or the non-secure domain. In theevent that the core is running in the secure domain, the processbranches to step 302, where a lookup is performed within the micro-TLB206 to see if the relevant portion of the virtual address matches withone of the virtual address portions within the micro-TLB. In the eventof a hit at step 302, the process branches directly to step 312, wherethe access permission logic 202 performs the necessary access permissionanalysis. At step 314, it is then determined whether there is an accesspermission violation, and if there is the process proceeds to step 316,where the access permission logic 202 issues an abort over path 230.Otherwise, in the absence of such an access permission violation, theprocess proceeds from step 314 to step 318, where the memory accessproceeds. In particular the region attribute logic 204 will output thenecessary cacheable and bufferable attributes over path 232, and themicro-TLB 206 will issue the physical address over path 238 as describedearlier.

[0453] If at step 302 there is a miss in the micro-TLB, then a lookupprocess is performed within the main TLB 208 at step 304 to determinewhether the required secure descriptor is present within the main TLB.If not, then a page table walk process is executed at step 306, wherebythe translation table walk logic 210 obtains the required descriptorfrom the secure page table, as described earlier with reference to FIG.37. The process then proceeds to step 308, or proceeds directly to step308 from step 304 in the event that the secure descriptor was already inthe main TLB 208.

[0454] At step 308, it is determined that the main TLB now contains thevalid tagged secure descriptor, and accordingly the process proceeds tostep 310, where the micro-TLB is loaded with the sub-section of thedescriptor that contains the physical address portion. Since the core 10is currently running in secure mode, there is no need for the partitionchecker 222 to perform any partition checking function.

[0455] The process then proceeds to step 312 where the remainder of thememory access proceeds as described earlier.

[0456] In the event of a non-secure memory access, the process proceedsfrom step 300 to step 320, where a lookup process is performed in themicro-TLB 206 to determine whether the corresponding physical addressportion from a non-secure descriptor is present. If it is, then theprocess branches directly to step 336, where the access permissionrights are checked by the access permission logic 202. It is importantto note at this point that if the relevant physical address portion iswithin the micro-TLB, it is assumed that there is no security violation,since the partition checker 222 effectively polices the informationprior to it being stored within the micro-TLB, such that if theinformation is within the micro-TLB, it is assumed to be the appropriatenon-secure information. Once the access permission has been checked atstep 336, the process proceeds to step 338, where it is determinedwhether there is any violation, in which event an access permissionfault abort is issued at step 316. Otherwise, the process proceeds tostep 318 where the remainder of the memory access is performed, asdiscussed earlier.

[0457] In the event that at step 320 no hit was located in themicro-TLB, the process proceeds to step 322, where a lookup process isperformed in the main TLB 208 to determine whether the relevantnon-secure descriptor is present. If not, a page table walk process isperformed at step 324 by the translation table walk logic 210 in orderto retrieve into the main TLB 208 the necessary non-secure descriptorfrom the non-secure page table. The process then proceeds to step 326,or proceeds directly to step 326 from step 322 in the event that a hitwithin the main TLB 208 occurred at step 322. At step 326, it isdetermined that the main TLB now contains the valid tagged non-securedescriptor for the virtual address in question, and then at step 328 thepartition checker 222 checks that the physical address that would begenerated from the virtual address of the memory access request (giventhe physical address portion within the descriptor) will point to alocation in non-secure memory. If not, i.e. if the physical addresspoints to a location in secure memory, then at step 330 it is determinedthat there is a security violation, and the process proceeds to step 332where a secure/non-secure fault abort is issued by the partition checker222.

[0458] If however the partition checker logic 222 determines that thereis no security violation, the process proceeds to step 334, where themicro-TLB is loaded with the sub-section of the relevant descriptor thatcontains the physical address portion, whereafter at step 336 the memoryaccess is then processed in the earlier described manner.

[0459] The handling of memory access requests that directly issue aphysical address will now be described with reference to FIG. 40. Asmentioned earlier, in this scenario, the MMU 200 will be deactivated,this preferably being achieved by the setting within a relevant registerof the CP15 registers an MMU enable bit, this setting process beingperformed by the monitor mode. Hence, at step 350 the core 10 willgenerate a physical address which will be passed over path 236 into theMPU 220. Then, at step 352, the MPU checks permissions to verify thatthe memory access being requested can proceed given the current mode ofoperation, i.e. user, supervisor, etc. Furthermore, if the core isoperating in non-secure mode, the partition checker 222 will also checkat step 352 whether the physical address is within non-secure memory.Then, at step 354, it is determined whether there is a violation, i.e.whether the access permission processing has revealed a violation, or ifin non-secure mode, the partition checking process has identified aviolation. If either of these violations occurs, then the processproceeds to step 356 where an access permission fault abort is generatedby the MPU 220. It will be appreciated that in certain embodiments theremay be no distinction between the two types of abort, whereas inalternative embodiments the abort signal could indicate whether itrelates to an access permission fault or a security fault.

[0460] If no violation is detected at step 354, the process proceeds tostep 358, where the memory access to the location identified by thephysical address occurs.

[0461] In preferred embodiments only the monitor mode is arranged togenerate physical addresses directly, and accordingly in all other casesthe MMU 200 will be active and generation of the physical address fromthe virtual address of the memory access request will occur as describedearlier.

[0462]FIG. 38 illustrates an alternative embodiment of the memorymanagement logic in a situation where all memory access requests specifya virtual address, and accordingly physical addresses are not generateddirectly in any of the modes of operation. In this scenario, it will beappreciated that a separate MPU 220 is not required, and instead thepartition checker 222 can be incorporated within the MMU 200. Thischange aside, the processing proceeds in exactly the same manner asdiscussed earlier with reference to FIGS. 37 and 39.

[0463] It will be appreciated that various other options are alsopossible. For example, assuming memory access requests may be issued byboth secure and non-secure modes specifying virtual addresses, two MMUscould be provided, one for secure access requests and one for non-secureaccess requests, i.e. MPU 220 in FIG. 37 could be replaced by a completeMMU. In such cases, the use of flags with the main TLB of each MMU todefine whether descriptors are secure or non-secure would not be needed,as one MMU would store non-secure descriptors in its main TLB, and theother MMU would store secure descriptors in its main TLB. Of course, thepartition checker would still be required to check whether an access tosecure memory is being attempted whilst the core is in the non-securedomain.

[0464] If, alternatively, all memory access requests directly specifiedphysical addresses, an alternative implementation might be to use twoMPUs, one for secure access requests and one for non-secure accessrequests. The MPU used for non-secure access requests would have itsaccess requests policed by a partition checker to ensure accesses tosecure memory are not allowed in non-secure modes.

[0465] As a further feature which may be provided with either the FIG.37 or the FIG. 38 arrangement, the partition checker 222 could bearranged to perform some partition checking in order to police theactivities of the translation table walk logic 210. In particular, ifthe core is currently operating in the non-secure domain, then thepartition checker 222 could be arranged to check, whenever thetranslation table walk logic 210 is seeking to access a page table, thatit is accessing the non-secure page table rather than the secure pagetable. If a violation is detected, an abort signal would preferably begenerated. Since the translation table walk logic 210 typically performsthe page table lookup by combining a page table base address withcertain bits of the virtual address issued by the memory access request,this partition checking may involve, for example, checking that thetranslation table walk logic 210 is using a base address of a non-securepage table rather than a base address of a secure page table.

[0466]FIG. 41 illustrates schematically the process performed by thepartition checker 222 when the core 10 is operating in a non-securemode. It will be appreciated that in normal operation a descriptorobtained from the non-secure page table should describe a page mapped innon-secure memory only. However, in the case of software attack, thedescriptor may be tampered with in order that it now describes a sectionthat contains both non-secure and secure regions of memory. Hence,considering the example in FIG. 41, the corrupted non-secure descriptormay cover a page that includes non-secure areas 370, 372, 374 and secureareas 376, 378, 380. If the virtual address issued as part of the memoryaccess request would then correspond to a physical address in a securememory region, for example the secure memory region 376 as illustratedin FIG. 41, then the partition checker 222 is arranged to generate anabort to prevent that access taking place. Hence, even though thenon-secure descriptor has been corrupted in an attempt to gain access tosecure memory, the partition checker 222 prevents the access takingplace. In contrast, if the physical address that would be derived usingthis descriptor corresponds to a non-secure memory region, for exampleregion 374 as illustrated in FIG. 41, then the access controlinformation loaded into the micro-TLB 206 merely identifies thisnon-secure region 374. Hence, accesses within that non-secure memoryregion 374 can occur but no accesses into any of the secure regions 376,378 or 380 can occur. Thus, it can be seen that even though the main TLB208 may contain descriptors from the non-secure page table that havebeen tampered with, the micro-TLB will only contain physical addressportions that will enable access to non-secure memory regions.

[0467] As described earlier, in embodiments where both non-secure modesand secure modes may generate memory access requests specifying virtualaddresses, then the memory preferably comprises both a non-secure pagetable within non-secure memory, and a secure page table within securememory. When in non-secure mode, the non-secure page table will bereferenced by the translation table walk logic 210, whereas when insecure mode, the secure page table will be referenced by the translationtable walk logic 210. FIG. 42 illustrates these two page tables. Asshown in FIG. 42, the non-secure memory 390, which may for example bewithin external memory 56 of FIG. 1, includes within it a non-securepage table 395 specified in a CP15 register 34 by reference to a baseaddress 397. Similarly, within secure memory 400, which again may bewithin the external memory 56 of FIG. 1, a corresponding secure pagetable 405 is provided which is specified within a duplicate CP15register 34 by a secure page table base address 407. Each descriptorwithin the non-secure page table 395 will point to a correspondingnon-secure page in non-secure memory 390, whereas each descriptor withinthe secure page table 405 will define a corresponding secure page in thesecure memory 400. In addition, as will be described in more detaillater, it is possible for certain areas of memory to be shared memoryregions 410, which are accessible by both non-secure modes and securemodes.

[0468]FIG. 43 illustrates in more detail the lookup process performedwithin the main TLB 208 in accordance with preferred embodiments. Asmentioned earlier, the main TLB 208 includes a domain flag 425 whichidentifies whether the corresponding descriptor 435 is from the securepage table or the non-secure page table. This ensures that when a lookupprocess is performed, only the descriptors relevant to the particulardomain in which the core 10 is operating will be checked. FIG. 43illustrates an example where the core is running in the secure domain,also referred to as the secure world. As can be seen from FIG. 43, whena main TLB 208 lookup is performed, this will result in the descriptors440 being ignored, and only the descriptors 445 being identified ascandidates for the lookup process.

[0469] In accordance with preferred embodiments, an additional processID flag 430, also referred to herein as the ASID flag, is provided toidentify descriptors from process specific page tables. Accordingly,processes P1, P2 and P3 may each have corresponding page tables providedwithin the memory, and further may have different page tables fornon-secure operation and secure operation. Further, it will beappreciated that the processes P1, P2, P3 in the secure domain may beentirely separate processes to the processes P1, P2, P3 in thenon-secure domain. Accordingly, as shown in FIG. 43, in addition tochecking the domain when a main TLB lookup 208 is required, the ASIDflag is also checked.

[0470] Accordingly, in the example in FIG. 43 where in the securedomain, process P1 is executing, this lookup process identifies just thetwo entries 450 within the main TLB 208, and a hit or miss is thengenerated dependent on whether the virtual address portion within thosetwo descriptors matches with the corresponding portion of the virtualaddress issued by the memory access request. If it does, then therelevant access control information is extracted and passed to themicro-TLB 206, the access permission logic 202 and the region attributelogic 204. Otherwise, a miss occurs, and the translation table walklogic 210 is used to retrieved into the main TLB 208 the requireddescriptor from the page table provided for secure process P1. As willbe appreciated by those skilled in the art, there are many techniquesfor managing the content of a TLB, and accordingly when a new descriptoris retrieved for storage in the main TLB 208, and the main TLB isalready full, any one of a number of known techniques may be used todetermine which descriptor to evict from the main TLB to make room forthe new descriptor, for example least recently used approaches, etc.

[0471] It will be appreciated that the secure kernel used in securemodes of operation may be developed entirely separately to thenon-secure operating system. However, in certain cases the secure kerneland the non-secure operating system development may be closely linked,and in such situations it may be appropriate to allow secureapplications to use the non-secure descriptors. Indeed, this will allowthe secure applications to have direct access to non-secure data (forsharing) by knowing only the virtual address. This of course presumesthat the secure virtual mapping and the non-secure virtual mapping areexclusive for a particular ASID. In such scenarios, the tag introducedpreviously (i.e. the domain flag) to distinguish between secure andnon-secure descriptors will not be needed. The lookup in the TLB isinstead then performed with all of descriptors available.

[0472] In preferred embodiments, the choice between this configurationof the main TLB, and the earlier described configuration with separatesecure and non-secure descriptors, can be set by a particular bitprovided within the CP15 control registers. In preferred embodiments,this bit would only be set by the secure kernel.

[0473] In embodiments where the secure application were directly allowedto use a non-secure virtual address, it would be possible to make anon-secure stack pointer available from the secure domain. This can bedone by copying a non-secure register value identifying the non-securestack pointer into a dedicated register within the CP15 registers 34.This will then enable the non-secure application to pass parameters viathe stack according to a scheme understood by the secure application.

[0474] As described earlier, the memory may be partitioned intonon-secure and secure parts, and this partitioning is controlled by thesecure kernel using the CP15 registers 34 dedicated to the partitionchecker 222. The basic partitioning approach is based on region accesspermissions as definable in typical MPU devices. Accordingly, the memoryis divided into regions, and each region is preferably defined with itsbase address, size, memory attributes and access permissions. Further,when overlapping regions are programmed, the attributes of the upperregion take highest priority. Additionally, in accordance with preferredembodiments of the present invention, a new region attribute is providedto define whether that corresponding region is in secure memory or innon-secure memory. This new region attribute is used by the securekernel to define the part of the memory that is to be protected assecure memory.

[0475] At the boot stage, a first partition is performed as illustratedin FIG. 44. This initial partition will determine the amount of memory460 allocated to the non-secure world, non-secure operating system andnon-secure applications. This amount corresponds to the non-secureregion defined in the partition. This information will then be used bythe non-secure operating system for its memory management. The rest ofthe memory 462, 464, which is defined as secure, is unknown by thenon-secure operating system. In order to protect integrity in thenon-secure world, the non-secure memory may be programmed with accesspermission for secure privileged modes only. Hence, secure applicationswill not corrupt the non-secure ones. As can be seen from FIG. 44,following this boot stage partition, memory 460 is available for use bythe non-secure operating system, memory 462 is available for use by thesecure kernel, and memory 464 is available for use by secureapplications.

[0476] Once the boot stage partition has been performed, memory mappingof the non-secure memory 460 is handled by the non-secure operatingsystem using the MMU 200, and accordingly a series of non-secure pagescan be defined in the usual manner. This is illustrated in FIG. 45.

[0477] If a secure application needs to share memory with a non-secureapplication, the secure kernel can change the rights of a part of thememory to transfer artificially data from one domain to the other.Hence, as illustrated in FIG. 46, the secure kernel can, after checkingthe integrity of a non-secure page, change the rights of that page suchthat it becomes a secure page 466 accessible as shared memory.

[0478] When the partition of the memory is changed, the micro-TLB 206needs to be flushed. Hence, in this scenario, when a non-secure accesssubsequently occurs, a miss will occur in the micro-TLB 206, andaccordingly a new descriptor will be loaded from the main TLB 208. Thisnew descriptor will subsequently be checked by the partition checker 222of the MPU as it is attempted to retrieve it into the micro-TLB 206, andso will be consistent with the new partition of the memory.

[0479] In preferred embodiments, the cache 38 is virtual-indexed andphysical-tagged. Accordingly, when an access is performed in the cache38, a lookup will have already been performed in the micro-TLB 206first, and accordingly access permissions, especially secure andnon-secure permissions, will have been checked. Accordingly, secure datacannot be stored in the cache 38 by non-secure applications. Access tothe cache 38 is under the control of the partition checking performed bythe partition checker 222, and accordingly no access to secure data canbe performed in non-secure mode.

[0480] However, one problem that could occur would be for an applicationin the non-secure domain to be able to use the cache operations registerto invalidate, clean, or flush the cache. It needs to be ensured thatsuch operations could not affect the security of the system. Forexample, if the non-secure operating system were to invalidate the cache38 without cleaning it, any secure dirty data must be written to theexternal memory before being replaced. Preferably, secure data is taggedin the cache, and accordingly can be dealt with differently if desired.

[0481] In preferred embodiments, if an “invalidate line by address”operation is executed by a non-secure program, the physical address ischecked by the partition checker 222, and if the cache line is a securecache line, the operation becomes a “clean and invalidate” operation,thereby ensuring that the security of the system is maintained. Further,in preferred embodiments, all “invalidate line by index” operations thatare executed by a non-secure program become “clean and invalidate byindex” operations. Similarly, all “invalidate all” operations executedby a non-secure program become “clean and invalidate all” operations.

[0482] Furthermore, with reference to FIG. 1, any access to the TCM 36by the DMA 32 is controlled by the micro-TLB 206. Hence, when the DMA 32performs a lookup in the TLB to translate its virtual address into aphysical one, the earlier described flags that were added in the mainTLB allow the required security checking to be performed, just as if theaccess request had been issued by the core 10. Further, as will bediscussed later, a replica partition checker is coupled to the externalbus 70, preferably being located within the arbiter/decoder block 54,such that if the DMA 32 directly accesses the memory coupled to theexternal bus 70 via the external bus interface 42, the replica partitionchecker connected to that external bus checks the validity of theaccess. Furthermore, in certain preferred embodiments, it would bepossible to add a bit to the CP15 registers 34 to define whether the DMAcontroller 32 can be used in the non-secure domain, this bit only beingallowed to be set by the secure kernel when operating in a privilegedmode.

[0483] Considering the TCM 36, if secure data is to placed within theTCM 36, this must be handled with care. As an example, a scenario couldbe imagined where the non-secure operating system programs the physicaladdress range for the TCM memory 36 so that it overlaps an externalsecure memory part. If the mode of operation then changes to a securemode, the secure kernel may cause data to be stored in that overlappingpart, and typically the data would be stored in the TCM 36, since theTCM 36 will typically have a higher priority than the external memory.If the non-secure operating system were then to change the setting ofthe physical address space for the TCM 36 so that the previous secureregion is now mapped in a non-secure physical area of memory, it will beappreciated that the non-secure operating system can then access thesecure data, since the partition checker will see the area as non-secureand won't assert an abort. Hence, to summarise, if the TCM is configuredto act as normal local RAM and not as SmartCache, it may be possible forthe non-secure operating system to read secure world data if it can movethe TCM base register to non-secure physical address.

[0484] To prevent this kind of scenario, a control bit is in preferredembodiments provided within the CP15 registers 34 which is onlyaccessible in secure privilege modes of operation, and provides twopossible configurations. In a first configuration, this control bit isset to “1”, in which event the TCM can only be controlled by the secureprivilege modes. Hence, any non-secure access attempted to the TCMcontrol registers within the CP15 34 will cause an undefined instructionexception to be entered. Thus, in this first configuration, both securemodes and non-secure modes can use the TCM, but the TCM is controlledonly by the secure privilege mode. In the second configuration, thecontrol bit is set to “0”, in which event the TCM can be controlled bythe non-secure operating system. In this case, the TCM is only used bythe non-secure applications. No secure data can be stored to or loadedfrom the TCM. Hence, when a secure access is performed, no look-up isperformed within the TCM to see if the address matched the TCM addressrange.

[0485] By default, it is envisaged that the TCM would be used only bynon-secure operating systems, as in this scenario the non-secureoperating system would not need to be changed.

[0486] As mentioned earlier, in addition to the provision of thepartition checker 222 within the MPU 220, preferred embodiments of thepresent invention also provide an analogous partition checking blockcoupled to the external bus 70, this additional partition checker beingused to police accesses to memory by other master devices, for examplethe digital signal processor (DSP) 50, the DMA controller 52 coupleddirectly to the external bus, the DMA controller 32 connectable to theexternal bus via the external bus interface 42, etc. Indeed in someembodiments, as will be discussed later, it is possible to solely have apartition checking block coupled to the external (or device) bus, andnot to provide a partition checker as part of the memory managementlogic 30. In some such embodiments, a partition checker may optionallybe provided as part of the memory management logic 30, in such instancesthis partition checker be considered as a further partition checkerprovided in addition to the one coupled to the device bus.

[0487] As mentioned earlier, the entire memory system can consist ofseveral memory units, and a variety of these may exist on the externalbus 70, for example the external memory 56, boot ROM 44, or indeedbuffers or registers 48, 62, 66 within peripheral devices such as thescreen driver 46, I/O interface 60, key storage unit 64, etc.Furthermore, different parts of the memory system may need to be definedas secure memory, for example it may be desired that the key buffer 66within the key storage unit 64 should be treated as secure memory. If anaccess to such secure memory were to be attempted by a device coupled tothe external bus, then it is clear that the earlier described memorymanagement logic 30 provided within the chip containing the core 10would not be able to police such accesses.

[0488]FIG. 47 illustrates how the additional partition checker 492coupled to the external bus, also referred to herein as the device bus,is used. The external bus would typically be arranged such that whenevermemory access requests were issued onto that external bus by devices,such as devices 470, 472, those memory access requests would alsoinclude certain signals on the external bus defining the mode ofoperation, for example privileged, user, etc. In accordance withpreferred embodiments of the present invention the memory access requestalso involves issuance of a domain signal onto the external bus toidentify whether the device is operating in secure mode or non-securemode. This domain signal is preferably issued at the hardware level, andin preferred embodiments a device capable of operating in secure ornon-secure domains will include a predetermined pin for outputting thedomain signal onto path 490 within the external bus. For the purpose ofillustration, this path 490 is shown separately to the other signalpaths 488 on the external bus.

[0489] This domain signal, also referred to herein as the “S bit” willidentify whether the device issuing the memory access request isoperating in secure domain or non-secure domain, and this informationwill be received by the partition checker 492 coupled to the externalbus. The partition checker 492 will also have access to the partitioninginformation identifying which regions of memory are secure ornon-secure, and accordingly can be arranged to only allow a device tohave access to a secure part of memory if the S bit is asserted toidentify a secure mode of operation.

[0490] By default, it is envisaged that the S bit would be unasserted,and accordingly a pre-existing non-secure device, such as device 472illustrated in FIG. 47, would not output an asserted S bit andaccordingly would never be granted access by the partition checker 492to any secure parts of memory, whether that be within registers orbuffers 482, 486 within the screen driver 480, the I/O interface 484, orwithin the external memory 474.

[0491] For the sake of illustration, the arbiter block 476 used toarbitrate between memory access requests issued by master devices, suchas devices 470, 472, is illustrated separately to the decoder 478 usedto determine the appropriate memory device to service the memory accessrequest, and separate from the partition checker 492. However, it willbe appreciated that one or more of these components may be integratedwithin the same unit if desired.

[0492]FIG. 48 illustrates an alternative embodiment, in which apartition checker 492 is not provided, and instead each memory device474, 480, 484 is arranged to police its own memory access dependent onthe value of the S bit. Accordingly, if device 470 were to assert amemory access request in non-secure mode to a register 482 within thescreen driver 480 that was marked as secure memory, then the screendriver 480 would determine that the S bit was not asserted, and wouldnot process the memory access request. Accordingly, it is envisaged thatwith appropriate design of the various memory devices, it may bepossible to avoid the need for a partition checker 492 to be providedseparately on the external bus.

[0493] In the above description of FIGS. 47 and 48, the “S bit” is saidto identify whether the device issuing the memory access request isoperating in secure domain or non-secure domain. Viewed another way,this S bit can be seen to indicate whether the memory access requestpertains to the secure domain or the non-secure domain.

[0494] In the embodiments described with reference to FIGS. 37 and 38, asingle MMU, along with a single set of page tables, was used to performvirtual to physical address translation. With such an approach, thephysical address space would typically be segmented between non-securememory and secure memory in a simplistic manner such as illustrated inFIG. 49. Here a physical address space 2100 includes an address spacestarting at address zero and extending to address Y for one of thememory units within the memory system, for example the external memory56. For each memory unit, the addressable memory would typically besectioned into two parts, a first part 2110 being allocated asnon-secure memory and a second part 2120 being allocated as securememory.

[0495] With such an approach, it will be appreciated that there arecertain physical addresses which are not accessible to particulardomain(s), and these gaps would be apparent to the operating system usedin those domain(s). Whilst the operating system used in the securedomain will have knowledge of the non-secure domain, and hence will notbe concerned by this, the operating system in the non-secure domainshould ideally not need to have any knowledge of the presence of thesecure domain, but instead should operate as though the secure domainwere not there.

[0496] As a further issue, it will be appreciated that a non-secureoperating system will see its address space for the external memory asstarting at address zero and extending to address X, and the non-secureoperating system need know nothing about the secure kernel and inparticular the presence of the secure memory extending from address X+1up to address Y. In contrast, the secure kernel will not see its addressspace beginning at address zero, which is not what an operating systemwould typically expect.

[0497] One embodiment which alleviates the above concerns by allowingthe secure memory regions to be completely hidden from the non-secureoperating system's view of its physical address space, and by enablingboth the secure kernel in the secure domain and the non-secure operatingsystem in the non-secure domain to see their address space for externalmemory as beginning at address zero is illustrated schematically in FIG.51. Here, the physical address space 2200 is able to be segmented at thepage level into either secure or non-secure segments. In the exampleillustrated in FIG. 51, the address space for the external memory isshown as being segmented into four sections 2210, 2220, 2230, and 2240,consisting of two secure memory regions and two non-secure memoryregions.

[0498] Rather than transitioning between the virtual address space andthe physical address space via a single page table conversion, twoseparate layers of address translation are performed with reference to afirst page table and a second page table, thereby enabling the conceptof an intermediate address space to be introduced which can be arrangeddifferently, dependent on whether the processor is in the secure domainor the non-secure domain. More particularly, as illustrated in FIG. 51,the two secure memory regions 2210 and 2230 in the physical addressspace can be mapped to the single region 2265 in the intermediateaddress space for the secure domain by use of descriptors providedwithin a secure page table within the set of page tables 2250. As far asthe operating system running on the processor is concerned, it will seethe intermediate address space as being the physical address space, andwill use an MMU to convert virtual addresses into intermediate addresseswithin the intermediate address space.

[0499] Similarly, an intermediate address space 2270 can be configuredfor the non-secure domain, in which the two non-secure memory regions2220 and 2240 in the physical address space are mapped to the non-secureregion 2275 in the intermediate address space for the non-secure domainvia corresponding descriptors in a non-secure page table within the setof page tables 2250.

[0500] In one embodiment, the translation of virtual addresses intophysical addresses via intermediate addresses is handled using twoseparate MMUs as illustrated in FIG. 50A. Each of the MMUs 2150 and 2170in FIG. 50A can be considered as being constructed in a similar mannerto the MMu 200 shown in FIG. 37, but for the sake of ease ofillustration certain detail has been omitted in FIG. 50A.

[0501] The first MMU 2150 includes a micro-TLB 2155, a main TLB 2160 andtranslation table walk logic 2165, while similarly the second MMU 2170includes a micro-TLB 2175, a main TLB 2180 and translation table walklogic 2185. The first MMU may be controlled by the non-secure operatingsystem when the processor is operating in the non-secure domain, or bythe secure kernel when the processor is operating in the secure domain.However, in preferred embodiments, the second MMU is only controllableby the secure kernel, or by the monitor program.

[0502] When the processor core 10 issues a memory access request, itwill issue a virtual address over path 2153 to the micro-TLB 2155. Themicro-TLB 2155 will store for a number of virtual address portionscorresponding intermediate address portions retrieved from descriptorsstored within the main TLB 2160, the descriptors in the main TLB 2160having been retrieved from page tables in a first set of page tablesassociated with the first MMU 2150. If a hit is detected within themicro-TLB 2155, then the micro-TLB 2155 can issue over path 2157 anintermediate address corresponding to the virtual address received overpath 2153. If there is no hit within the micro-TLB 2155, then the mainTLB 2160 will be referenced to see if a hit is detected within the mainTLB, and if so the relevant virtual address portion and correspondingintermediate address portion will be retrieved into the micro-TLB 2155,whereafter the intermediate address can be issued over path 2157.

[0503] If there is no hit within the micro-TLB 2155 and the main TLB2160, then the translation table walk logic 2165 is used to issue arequest for the required descriptor from a predetermined page table in afirst set of page tables accessible by the first MMU 2150. Typically,there may be page tables associated with individual processes for bothsecure domain or non-secure domain, and the intermediate base addressesfor those page tables will be accessible by the translation table walklogic 2165, for example from appropriate registers within the CP15registers 34. Accordingly, the translation table walk logic 2165 canissue an intermediate address over path 2167 to request a descriptorfrom the appropriate page table.

[0504] The second MMU 2170 is arranged to receive any intermediateaddresses output by the micro-TLB 2155 over path 2157, or by thetranslation table walk logic 2165 over path 2167, and if a hit isdetected within the micro-TLB 2175, the micro-TLB can then issue therequired physical address over path 2192 to memory to cause the requireddata to be retrieved over the data bus 2190. In the event of anintermediate address issued over path 2157, this will cause the requireddata to be returned to the core 10, whilst for an intermediate addressissued over path 2167, this will cause the required descriptor to bereturned to the first MMU 2150 for storage within the main TLB 2160.

[0505] In the event of a miss in the micro-TLB 2175, the main TLB 2180will be referenced, and if there is a hit within the main TLB, therequired intermediate address portion and corresponding physical addressportion will be returned to the micro-TLB 2175, to then enable themicro-TLB 2175 to issue the required physical address over path 2192.However, in the absence of a hit in either the micro-TLB 2175 or themain TLB 2180, then the translation table walk logic 2185 will bearranged to output a request over path 2194 for the required descriptorfrom the relevant page table within a second set of page tablesassociated with the second MMU 2170. This second set of page tablesincludes descriptors which associate intermediate address portions withphysical address portions, and typically there will be at least one pagetable for secure domain and one page table for non-secure domain. When arequest is issued over path 2194, this will result in the relevantdescriptor from the second set of page tables being returned to thesecond MMU 2170 for storing within the main TLB 2180.

[0506] The operation of the embodiment illustrated in FIG. 50A will nowbe illustrated further by way of a specific example as set out below, inwhich the abbreviation VA denotes virtual address, IA denotesintermediate address, and PA denotes physical address: 1) Core issues VA= 3000 [IA = 5000, PA = 7000] 2) Miss in micro-TLB of MMU 1 3) Miss inmain TLB of MMU 1 Page Table 1 Base Address = 8000 IA [PA = 10000] 4)Translation Table Walk logic in MMU 1 performs page table lookup -issues IA = 8003 5) Miss in micro-TLB of MMU 2 6) Miss in main TLB ofMMU 2 Page Table 2 Base Address = 12000 PA 7) Translation Table WalkLogic in MMU 2 performs page table lookup - issues PA = 12008 “8000 IA =10000 PA” returned as page table data 8) - stored in main TLB of MMU 29) - stored in micro-TLB of MMU 2 10) Micro-TLB in MMU 2 now has hit -issues PA = 10003 “3000 VA = 5000 IA” returned as page table data 11) -stored in main TLB of MMU 1 12) - stored in micro-TLB of MMU 1 13)Micro-TLB in MMU 1 now has hit issues IA = 5000 to perform data access14) miss in micro-TLB of MMU 2 15) miss in main TLB of MMU 2 16)Translation Table Walk Logic in MMU 2 performs page table lookup -issues PA = 12005 “5000 IA = 7000 PA” returned as page table data 17) -stored in main TLB of MMU 2 18) - stored in micro-TLB of MMU 2 19)Micro-TLB in MMU 2 now has hit - issues PA = 7000 to perform data access20) Data at physical address 7000 returned to core

NEXT TIME CORE ISSUES A MEMORY ACCESS REQUEST (say VA 3001 . . . )

[0507] 1) Core issues VA=3001

[0508] 2) Hit in micro-TLB of MMU1, request IA 5001 issued to MMU2

[0509] 3) Hit in micro-TLB on MMU2, request for PA 7001 issued to memory

[0510] 4) Data at PA 7001 returned to core.

[0511] It will be appreciated that in the above example misses occur inboth the micro-TLB and the main TLB of both MMUs, and hence this examplerepresents the ‘worst case’ scenario. Typically, it would be expectedthat a hit would be observed in at least one of the micro-TLBs or mainTLBs, thereby significantly reducing the time taken to retrieve thedata.

[0512] Returning to FIG. 51, the second set of page tables 2250 willtypically be provided within a certain region of the physical addressspace, in preferred embodiments a secure region. The first set of pagetables can be split into two types, namely secure page tables andnon-secure page tables. Preferably, the secure page tables will appearconsecutively within the intermediate address space 2265, as will thenon-secure page tables within the non-secure intermediate address space2275. However, they need not be placed consecutively within the physicaladdress space, and accordingly, by way of example, the secure pagetables for the first set of page tables may be spread throughout thesecure regions 2210, 2230, and in a similar way the non-secure pagetables may be spread throughout the non-secure memory regions 2220 and2240.

[0513] As mentioned previously, one of the main benefits of using thetwo-level approach of two sets of page tables is that for both theoperating system of the secure domain and the operating system of thenon-secure domain the physical address space can be arranged to start atzero, which is what would typically be expected by an operating system.Additionally the secure memory regions can be completely hidden from thenon-secure operating system's view of its “physical address” space,since it sees as its physical address space the intermediate addressspace, which can be arranged to have a contiguous sequence ofintermediate addresses.

[0514] Additionally, the use of such an approach considerably simplifiesthe process of swapping regions of memory between non-secure memory andsecure memory. This is illustrated schematically with reference to FIG.52. As can be seen in FIG. 52, a region of memory 2300, which may forexample be a single page of memory, may exist within the non-securememory region 2220, and similarly a memory region 2310 may exist withinthe secure memory region 2210. However, these two memory regions 2300and 2310 can readily be swapped merely by changing the relevantdescriptors within the second set of page tables, such that the region2300 now becomes a secure region mapped to region 2305 in theintermediate address space of the secure domain, whilst region 2310 nowbecomes a non-secure region mapped to the region 2315 in theintermediate address space of the non-secure domain. This can occurentirely transparently to the operating systems in both the securedomain the non-secure domain, since their view of the physical addressspace is actually the intermediate address space of the secure domain ornon-secure domain, respectively. Hence, this approach avoids anyredefinition of the physical address space within each operating system.

[0515] An alternative embodiment of the present invention where two MMUsare also used, but in a different arrangement to that of FIG. 50A, willnow be described with reference to FIG. 50B. As can be seen from acomparison of FIG. 50B with FIG. 50A, the arrangement is almostidentical, but in this embodiment the first MMU 2150 is arranged toperform virtual address to physical address translation and the secondMMU is arranged to perform intermediate address to physical addresstranslation. Hence, instead of the path 2157 from the micro-TLB 2155 inthe first MMU 2150 to the micro-TLB 2175 in the second MMu 2170 used inthe FIG. 50A embodiment, the micro-TLB 2155 in the first MMU is insteadarranged to output a physical address directly over path 2192, as shownin FIG. 50B. The operation of the embodiment illustrated in FIG. 50Bwill now be illustrated by way of the specific example as set out below,which details the processing of the same core memory access request asillustrated earlier for the FIG. 50A embodiment: 1)  Core issues VA =3000 [IA = 5000, PA = 7000] 2)  Miss in micro-TLB and main TLB of MMU 1Page Table 1 Base Address = 8000 IA [PA = 10000] 3)  Translation TableWalk logic in MMU1 performs page table lookup - issues IA = 8003 4)  IA8003 misses in micro-TLB and main TLB of MMU 2 Page Table 2 Base Address= 12000 PA 5)  Translation Table Walk logic in MMU2 performs page tablelookup - issues PA = 12008 “8000 IA == 10000 PA” returned as page tabledata 6)  “8000 IA = 10000 PA” mapping stored in Main and micro-TLB ofMMU2 7)  Micro-TLB in MMU2 can now translate the request from step (3)to PA 10003 and issues fetch “3000 VA = 5000 IA” returned as page tabledata NOTE: This translation is retained in temporary storage by MMU1,but not stored directly in any TLB. 8)  Translation table walk logic ofMMU1 now issues request to MMU2 for IA = 5000 9)  IA 5000 misses in uTLBand main TLB of MMU 2 10)  Translation Table Walk logic in MMU2 performspage table lookup - issues PA = 12005 “5000 IA = 7000 PA” returned aspage table data 11)  MMU2 stores “5000 IA = 7000 PA” in uTLB and mainTLB. This translation is also communicated to MMU1. 12a)  MMU2 issuesthe PA = 7000 access to memory 12b)  MMU1 combines the “3000 VA = 5000IA” and the “5000 IA = 7000 PA” descriptors to give a “3000 VA = 7000PA” descriptor, which is stored in the main TLB and micro-TLB of MMU 1.13)  Data at PA 7000 is returned to the core.

NEXT TIME CORE ISSUES A MEMORY ACCESS REQUEST (say VA 3001 . . . )

[0516] 1) Core issues VA=3001

[0517] 2) Hit in micro-TLB of MMU1, MMU1 issues request for PA=7001

[0518] 3) Data at PA 7001 is returned to the core.

[0519] As can bee seen from a comparison of the above example with thatprovided for FIG. 50A, the main differences here are in step 7 whereMMU1 does not store the first table descriptor directly, and in step 12b (12 a and 12 b can happen at the same time) where MMU1 also receivesthe IA->PA translation and does the combination and stores the combineddescriptor in its TLBs.

[0520] Hence, it can be seen that whilst this alternative embodimentstill uses the two sets of page tables to convert virtual addresses tophysical addresses, the fact that the micro-TLB 2155 and main TLB 2160store the direct virtual address to physical address translation avoidsthe need for lookups to be performed in both MMUs when a hit occurs ineither the micro-TLB 2155 or the main TLB 2160. In such cases the firstMMU can directly handle requests from the core without reference to thesecond MMU.

[0521] It will be appreciated that the second MMU 2170 could be arrangednot to include the micro-TLB 2175 and the main TLB 2180, in which casethe page table walk logic 2185 would be used for every request thatneeded handling by the second MMU. This would save on complexity andcost for the second MMU, and might be acceptable assuming the second MMUwas only needed relatively infrequently. Since the first MMU will needto be used for every request, it will typically be expedient to includethe micro-TLB 2155 and main TLB 2160 in the first MMU 2150 to improvespeed of operation of the first MMU.

[0522] It should be noted that pages in the page tables may vary insize, and it is hence possible that the descriptors for the two halvesof the translation relate to different sized pages. Typically, the MMU1pages will be smaller than the MMU2 pages but this is not necessarilythe case. For example:

[0523] Table 1 maps 4 Kb at 0×40003000 onto 0×00081000

[0524] Table 2 maps 1 Mb at 0×00000000 onto 0×02000000

[0525] Here, the smallest of the two sizes must be used for the combinedtranslation, so the combined descriptor is

[0526] 4 Kb at 0×40003000 onto 0×02081000

[0527] However, where data is being swapped between worlds (as discussedearlier with reference to FIG. 52) it is possible that the reverse istrue, for example:

[0528] Table 1 maps 1 Mb at 0×0000000 onto 0×00000000

[0529] Table 2 maps 4 Kb at 0×00042000 onto 0×02042000

[0530] Now, a lookup at address 0×c0042010 from the core gives themapping:

[0531] 4 Kb at 0×c0042000 onto 0×02042000

[0532] i.e. the smaller of the two sizes is always used for the combinedmapping.

[0533] Note that in the second case the process is less efficient, sincethe (1 Mb) descriptor in table 1 will be repeatedly looked up anddiscarded as different 4Kb areas are accessed. However, in a typicalsystem the table 2 descriptors will be larger (as in the first example)the majority of the time, which is more efficient (the 1 Mb mapping canbe recycled for other 4 Kb pages which point into the appropriatesection of IA space).

[0534] As an alternative to employing two separate MMUs as illustratedin FIGS. 50A and 50B, a single MMU can be used as illustrated in FIG.53, where upon a miss in the main TLB 2420, an exception is generated bythe MMU, which then causes software to be run within the core 10 toproduce a virtual to physical address translation based on a combinationof descriptors from the two different sets of page tables. Moreparticularly, as shown in FIG. 53, the core 10 is coupled to an MMU2400, which includes a micro-TLB 2410 and a main TLB 2420. When the core10 issues a memory access request, the virtual address is provided overpath 2430, and if a hit is observed in the micro-TLB, then thecorresponding physical address is output directly over path 2440,causing the data to be returned over path 2450 into the core 10.However, if there is a miss in the micro-TLB 2410, the main TLB 2420 isreferenced and if the relevant descriptor is contained within the mainTLB the associated virtual address portion and corresponding physicaladdress portion are retrieved into the micro-TLB 2410, whereafter thephysical address can be issued over path 2440. However, if the main TLBalso produces a miss, then an exception is generated over path 2422 tothe core. The process performed within the core from receipt of such anexception will now be described further with reference to FIG. 54.

[0535] As shown in FIG. 54, if a TLB miss exception is detected by thecore at step 2500, then the core enters the monitor mode at apredetermined vector for that exception at step 2510. This will thencause page table merging code to be run to perform the remainder of thesteps illustrated in FIG. 54.

[0536] More particularly, at step 2520, the virtual address that wasissued over path 2430, and that gave rise to the miss in both themicro-TLB 2410 and the main TLB 2420 (hereafter referred to as thefaulting virtual address) is retrieved, whereafter at step 2530 theintermediate address for the required first descriptor is determineddependent on the intermediate base address for the appropriate tablewithin the first set of tables. Once that intermediate address has beendetermined (typically by some predetermined combination of the virtualaddress with the intermediate base address), then the relevant tablewithin the second set of tables is referenced in order to obtain thecorresponding physical address for that first descriptor. Thereafter atstep 2550 the first descriptor can be fetched from memory in order toenable the intermediate address for the faulting virtual address to bedetermined.

[0537] Then, at step 2560, the second table is again referenced to finda second descriptor giving the physical address for the intermediateaddress of the faulting virtual address. Thereafter at step 2570, thesecond descriptor is fetched to obtain the physical address for thefaulting virtual address.

[0538] Once the above information has been obtained, then the programmerges the first and second descriptors to generate a new descriptorgiving the required virtual address to physical address translation,this step being performed at step 2580. In a similar manner to thatdescribed earlier with reference to FIG. 50B, the merging performed bythe software again uses the smallest page table size for the combinedtranslation. Thereafter, at step 2590, the new descriptor is storedwithin the main TLB 2420, whereafter the process returns from theexception at step 2595.

[0539] Thereafter, the core 10 will be arranged to reissue the virtualaddress for the memory access request over path 2430, which will stillresult in a miss in the micro-TLB 2410, but will now result in a hit inthe main TLB 2420. Hence, the virtual address portion and correspondingphysical address portion can be retrieved into the micro-TLB 2410,whereafter the micro-TLB 2410 can then issue the physical address overpath 2440, resulting in the required data being returned to the core 10over path 2450.

[0540] It will be appreciated that, as alternative embodiments to thosedescribed earlier with reference to FIGS. 50A and 50B, one or both MMUsin those embodiments could be managed by software using the principlesdescribed above with reference to FIGS. 53 and 54.

[0541] Irrespective of whether two MMUs are used as shown in FIGS. 50Aor 50B, or one MMU is used as shown in FIG. 53, the fact that the secondset of page tables is managed by the processor when operating in monitormode (or alternatively in a privileged secure mode) ensures that thosepage tables are secure. As a result, when the processor is in thenon-secure domain it can only see non-secure memory, since it is onlythe intermediate address space generated for the non-secure domain bythe second set of page tables that the processor can see when in thenon-secure domain. As a result, there is no need to provide a partitionchecker as part of the memory management logic 30 illustrated in FIG. 1.However, a partition checker would still be provided on the external busto monitor accesses made by other bus masters in the system.

[0542] In the embodiments discussed earlier with reference to FIGS. 37and 38, a partition checker 222 was provided in association with the MMU200, and accordingly when an access is to be performed in the cache 38,a look-up will have already been performed in the micro-TLB 206 first,and accordingly access permissions, especially secure and non-securepermissions, would have been checked. Accordingly, in such embodiments,secure data cannot be stored in the cache 38 by non-secure applications.Access to the cache 38 is under the control of the partition checkingperformed by the partition checker 222, and accordingly no access tosecure data can be performed in non-secure mode.

[0543] However, in an alternative embodiment of the present invention, apartition checker 222 is not provided for monitoring accesses made overthe system bus 40, and instead the data processing apparatus merely hasa single partition checker coupled to the external bus 70 for monitoringaccesses to memory units connected to that external bus. In suchembodiments, this then means that the processor core 10 can access anymemory units coupled directly to the system bus 40, for example the TCM36 and the cache 38, without those accesses being policed by theexternal partition checker, and accordingly some mechanism is requiredto ensure that the processor core 10 does not access secure data withinthe cache 38 or the TCM 36 whilst operating in a non-secure mode.

[0544]FIG. 55 illustrates a data processing apparatus in accordance withone embodiment of the present invention, where a mechanism is providedto enable the cache 38 and/or the TCM 36 to control accesses made tothem without the need for any partition checking logic to be provided inassociation with the MMU 200. As shown in FIG. 55, the core 10 iscoupled via an MMU 200 to the system bus 40, the cache 38 and TCM 36also being coupled to the system bus 40. The core 10, cache 38 and TCM36 are coupled via the external bus interface 42 to the external bus 70,which as illustrated in FIG. 55 consists of an address bus 2620, acontrol bus 2630 and a data bus 2640.

[0545] The core 10, MMU 200, cache 38, TCM 36 and external bus interface42 can be viewed as constituting a single device connected onto theexternal bus 70, also referred to as a device bus, and other devices mayalso be coupled to that device bus, for example the secure peripheraldevice 470 or the non-secure peripheral device 472. Also connected tothe device bus 70 will be one or more memory units, for example theexternal memory 56. In addition, a bus control unit 2650 is connected tothe device bus 70, and will typically include an arbiter 2652, a decoder2654 and a partition checker 2656. For a general discussion of theoperation of the components connected to the device bus, referenceshould be made to the earlier described FIG. 47. In the earlierdescribed FIG. 47, the arbiter, decoder and partition checker were shownas separate blocks, but these elements work in the same manner whenplaced within a single control block 2650.

[0546] The MMU 200 of FIG. 55 is illustrated in more detail in FIG. 56.By comparison of FIG. 56 with FIG. 37, it can be seen that the MMU 200is constructed in exactly the same way as the MMU of FIG. 37, the onlydifference being that a partition checker 222 is not provided formonitoring data sent over path 242 between the main TLB 208 and themicro-TLB 206. If the processor core 10 issues a memory access requestthat specifies a virtual address, then that memory access request willbe routed through the MMU 200, and will be processed as describedearlier with reference to FIG. 37, resulting in a physical address beingoutput onto the system bus 40 over path 238 from the micro-TLB 206. If,in contrast, the memory access request directly specifies a physicaladdress, this will bypass the MMU 200, and instead will be routeddirectly onto the system bus 40 via path 236. In one embodiment onlywhen the processor is operating in the monitor mode will it generatememory access requests that directly specify physical addresses.

[0547] As will be recalled from the earlier description of the MMU 200,and in particular from the description of FIG. 43, the main TLB 208 willcontain a number of descriptors 435, and for each descriptor a domainflag 425 will be provided to identify whether the correspondingdescriptor is from a secure page table or a non-secure page table. Thesedescriptors 435 and associated domain flags 425 are illustratedschematically within the MMU 200 of FIG. 55.

[0548] When the core 10 issues a memory access request, this will resultin a physical address for that memory access request being output on thesystem bus 40 and typically the cache 38 will then perform a look-upprocess to determine whether the data item specified by that address isstored within the cache. Whenever a miss occurs within the cache, i.e.it is determined that the data item subject to the access request is notstored within the cache, a linefill procedure will be initiating by thecache in order to retrieve from the external memory 56 a line of datathat includes the data item the subject of the memory access request. Inparticular, the cache will output via the EBI 42 a linefill request ontothe control bus 2630 of the device bus 70, with a start address beingoutput on the address bus 2620. In addition, an HPROT signal will beoutput over path 2632 onto the control bus 2630, which will include adomain signal specifying the mode of operation of the core at the timethe memory access request was issued. Hence, the linefill process can beviewed as the propagation of the original memory access request onto theexternal bus by the cache 38.

[0549] This HPROT signal will be received by the partition checker 2656,and accordingly will identify to the partition checker whether thedevice requesting the specified data from the external memory 56 (inthis case the device incorporating the core 10 and the cache 38) wasoperating in the secure domain or the non-secure domain at the time thememory access request was issued. The partition checker 2656 will alsohave access to the partitioning information identifying which regions ofmemory are secure or non-secure, and accordingly can determine whetherthe device is allowed to have access to the data it is requesting.Hence, the partition checker can be arranged to only allow a device tohave access to a secure part of the memory if the domain signal withinthe HPROT signal (also referred to herein as an S bit) is asserted toidentify that access to this data was requested by the device whilstoperating in a secure mode of operation.

[0550] If the partition checker determines that the core 10 is notallowed to have access to the data requested, for example because theHPROT signal has identified that the core was operating in a non-securemode of operation, but the linefill request is seeking to retrieve datafrom the external memory that is within a secure region of memory, thenthe partition checker 2656 will issue an abort signal onto the controlbus 2630, which will be passed back over path 2636 to the EBI 42, andfrom there back to the cache 38, resulting in an abort signal beingissued over path 2670 to the core 10. However, if the partition checker2656 determines that the access is allowed, then it will output an S tagsignal identifying whether the data being retrieved from the externalmemory is secure data or non-secure data, and this S Tag signal will bepassed back via path 2634 to the EBI 42, and from there back to thecache 38 to enable setting of the flag 2602 associated with the cacheline 2600 the subject of the linefill process.

[0551] At the same time, the control logic 2650 will authorise theexternal memory 56 to output the linefill data requested, this databeing passed back via the EBI 42 over path 2680 to the cache 38 forstoring in the relevant cache line 2600. Hence, as a result of thisprocess, the chosen cache line within the cache 38 will be filled withdata items from the external memory 56, these data items including thedata item that was the subject of the original memory access requestfrom the core 10. The data item the subject of the memory access requestfrom the core can then be returned to the core from the cache 38, oralternatively can be provided directly from the EBI 42 back to the core10 over path 2660.

[0552] Since, in preferred embodiments, the original storage of data ina cache line will occur as a result of the above described linefillprocess, the flag 2602 associated with that cache line will be set basedon the value provided by the partition checker 2656, and that flag canthen be used by the cache 38 to directly control any subsequent accessto the data items in that cache line 2600. Hence, if the core 10subsequently issues a memory access request that produces a hit in aparticular cache line 2600 of the cache 38, the cache 38 will review thevalue of the associated flag 2602, and compare that value with thecurrent mode of operation of the core 10. In preferred embodiments, thiscurrent mode of operation of the core 10 is indicated by a domain bitset by the monitor mode within the CP15 domain status register. Hence,cache 38 can be arranged to only allow data items in a cache line thatthe corresponding flag 2602 indicates is secure data to be accessed bythe processor core 10 when the processor core 10 is operating in asecure mode of operation. Any attempt by the core to access secure datawithin the cache 38 whilst the core is operating in a non-secure modewill result in the cache 38 generating an abort signal over path 2670.

[0553] The TCM 36 can be set up in a variety of ways. In one embodiment,it can be set up to act like a cache, and in that embodiment will bearranged to include a plurality of lines 2610, each of which has a flag2612 associated therewith in the same way as the cache 38. Accesses tothe TCM 36 are then managed in exactly the same way as described earlierwith reference to the cache 38, with any TCM miss resulting in alinefill process being performed, as a result of which data will beretrieved into a particular line 2610, and the partition checker 2656will generate the required S tag value for storing in the flag 2612associated with that line 2610.

[0554] In an alternative embodiment, the TCM 36 may be set up as anextension of the external memory 56 and used to store data usedfrequently by the processor, since access to the TCM via the system busis significantly faster than access to external memory. In suchembodiments, the TCM 36 would not use the flags 2612, and instead adifferent mechanism would be used to control access to the TCM. Inparticular, as described earlier, in such embodiments, a control flagmay be provided which is settable by the processor when operating in aprivileged secure mode to indicate whether the tightly coupled memory iscontrollable by the processor only when executing in a privileged securemode or is controllable by the processor when executing in the at leastone non-secure mode. The control flag is set by the secure operatingsystem, and in effect defines whether the TCM is controlled by theprivileged secure mode or by non-secure modes. Hence, one configurationthat can be defmed is that the TCM is only controlled when the processoris operating in a privileged secure mode of operation. In suchembodiments, any non-secure access attempted to the TCM controlregisters will cause an undefined instruction exception to be entered.

[0555] In an alternative configuration, the TCM can be controlled by theprocessor when operating in a non-secure mode of operation. In suchembodiments, the TCM is only used by the non-secure applications. Nosecure data can be stored to or loaded from the TCM. Hence, when asecure access is performed, no look-up is performed within the TCM tosee if the address matched the TCM address range.

[0556]FIG. 57 is a flow diagram illustrating the processing performed bythe apparatus of FIG. 55 when a non-secure program operating on theprocessor core 10 generates a virtual address (step 2700). Firstly, atstep 2705, a look-up is performed within the micro-TLB 206, and if thisresults in a hit, the micro-TLB then checks access permissions at step2730. With reference to FIG. 56, this process can be viewed as beingperformed by the access permission logic 202.

[0557] If at step 2705, a miss occurs in the micro-TLB look-up, then alook-up is performed in the main TLB 208 amongst the non-securedescriptors stored therein (step 2710). If this results in a miss, thena page table walk process is performed at step 2715 (which has beendiscussed in detail previously with reference to FIG. 37), where afterat step 2720 it is determined that the main TLB contains the validtagged non-secure descriptor. If the look-up at step 2710 produces ahit, then the process proceeds directly to step 2720.

[0558] Thereafter, at step 2725, the micro-TLB is loaded with thesection of the descriptor which contains the physical address,whereafter at step 2730 the micro-TLB checks the access permissions.

[0559] If at step 2730, it is determined that there is a violation ofthe access permissions, then the process proceeds to step 2740, where anabort signal is issued over path 230 to the processor core (analogous topath 2670 shown in FIG. 55). However, assuming there is no violationdetected, then at step 2745 it is determined whether the access isrelated to a cacheable data item. If not, then an external access isinitiated at step 2790 to seek to retrieve the data item from theexternal memory 56. At step 2795, the partition checker 2656 willdetermine whether there is a secure partition violation, i.e. if theprocessor core 10 is seeking to access a data item in secure memorywhilst operating in a non-secure mode, and if a violation is detected,then the partition checker 2656 will generate an abort signal at step2775. However, assuming there is no secure partition violation, then theprocess proceeds to step 2785, where the data access takes place.

[0560] If at step 2745 it was determined that the data item beingrequested is cacheable, then a cache look-up is performed at step 2750within the cache, and if a hit is detected, the cache then determineswhether there is a secure line tag violation at step 2755. Hence, atthis stage, the cache will review the value of the flag 2602 associatedwith the cache line containing the data item, and will compare the valueof that flag with the mode of operation of the core 10 to determinewhether the core is entitled to access the data item requested. If asecure line tag violation is detected, then the process proceeds to step2760, where a secure violation fault abort signal is generated by thecache 38 and issued over path 2670 to the core 10. However, assumingthere is no secure line tag violation detected at step 2755, then thedata access is performed at step 2785.

[0561] If when the cache look-up is performed at step 2750 a cache missoccurs, then a cache linefill is initiated at step 2765. At step 2770,the partition checker 2656 then detects whether there is a securepartition violation, and if so issues an abort signal at step 2775.However, assuming there is no secure partition violation detected, thenthe cache linefill proceeds at step 2780, resulting in the data accesscompleting at step 2785.

[0562] As illustrated in FIG. 57, steps 2705, 2710, 2715, 2720, 2725,2730 and 2735 are performed within the MMU, steps 2745, 2750, 2755,2765, 2780 and 2790 are performed by the cache, and steps 2770 and steps2795 are performed by the partition checker.

[0563]FIG. 58 is a flow diagram showing the analogous process performedin the event that a secure program executing on the core generates avirtual address (step 2800). By comparison of FIG. 58 with FIG. 57, itwill be appreciated that steps 2805 through 2835 performed within theMMU are analogous to the steps 2705 through 2735 described earlier withreference to FIG. 57. The only difference is at step 2810, where thelook-up performed within the main TLB is performed in relation to anysecure descriptors stored within the main TLB, as a result of which atstep 2820 the main TLB contains valid tagged secure descriptors.

[0564] Within the cache, the cache no longer needs to look for anysecure line tag violation, since in the embodiment illustrated withreference to FIG. 58, it is assumed that the secure program can accessboth secure data and non-secure data. Accordingly, if a hit occursduring the cache look-up at step 2850, then the process proceedsdirectly to the data access step 2885.

[0565] Similarly, in the event that an external access to the externalmemory is required (i.e. at steps 2865 or 2890), the partition checkerneed perform no partition checking, since again it is assumed that thesecure program can access either secure data or non-secure data.

[0566] The steps 2845, 2850, 2865, 2880 and 2890 performed within thecache are analogous to the steps 2745, 2750, 2765, 2780 and 2790described earlier with reference to FIG. 57.

[0567]FIG. 59 shows different modes and applications running on aprocessor. The dashed lines indicate how different modes and/orapplications can be separated and isolated from one another duringmonitoring of the processor according to an embodiment of the presentinvention.

[0568] The ability to monitor a processor to locate possible faults anddiscover why an application is not performing as expected is extremelyuseful and many processors provide such functions. The monitoring can beperformed in a variety of ways including debug and trace functions.

[0569] In the processor according to the present technique debug canoperate in several modes including halt debug mode and monitor debugmode. These modes are intrusive and cause the program running at thetime to be stopped. In halt debug mode, when a breakpoint or watchpointoccurs, the core is stopped and isolated from the rest of the system andthe core enters debug state. On entry the core is halted, the pipelineis flushed and no instructions are pre-fetched. The PC is frozen and anyinterrupts (IRQ and FIQ) are ignored. It is then possible to examine thecore internal state (via the JTAG serial interface) as well as the stateof the memory system. This state is invasive to program execution, as itis possible to modify current mode, change register contents, etc. OnceDebug is terminated, the core exits from the Debug State by scanning inthe Restart instruction through the Debug TAP (test access port). Thenthe program resumes execution.

[0570] In monitor debug mode, a breakpoint or watchpoint causes the coreto enter abort mode, taking prefetch or Data Abort vectors respectively.In this case, the core is still in a functional mode and is not stoppedas it is in Halt debug mode. The abort handler communicates with adebugger application to access processor and coprocessor state or dumpmemory. A debug monitor program interfaces between the debug hardwareand the software debugger. If bit 11 of the debug status and controlregister DSCR is set (see later), interrupts (FIQ and IRQ) can beinhibited. In monitor debug mode, vector catching is disabled on DataAborts and Prefetch Aborts to avoid the processor being forced into anunrecoverable state as a result of the aborts that are generated for themonitor debug mode. It should be noted that monitor debug mode is a typeof debug mode and is not related to monitor mode of the processor whichis the mode that supervises switching between secure world andnon-secure world.

[0571] Debug can provide a snapshot of the state of a processor at acertain moment. It does this by noting the values in the variousregisters at the moment that a debug initiation request is received.These values are recorded on a scan chain (541, 544 of FIG. 67) and theyare then serially output using a JTAG controller (18 or FIG. 1).

[0572] An alternative way of monitoring the core is by trace. Trace isnot intrusive and records subsequent states as the core continues tooperate. Trace runs on an embedded trace macrocell (ETM) 22, 26 ofFIG. 1. The ETM has a trace port through which the trace information isexported, this is then analysed by an external trace port analyser.

[0573] The processor of embodiments of the present technique operates intwo separate domains, in the embodiments described these domainscomprise secure and non-secure domains. However, for the purposes of themonitoring functions, it will be clear to the skilled person that thesedomains can be any two domains between which data should not leak.Embodiments of the present technique are concerned with preventingleakage of data between the two domains and monitoring functions such asdebug and trace which are conventionally allowed access to the wholesystem are a potential source of data leakage between the domains.

[0574] In the example given above of a secure and non-secure domain orworld, secure data must not be available to the non-secure world.Furthermore, if debug is permitted, in secure world, it may beadvantageous for some of the data within secure world to be restrictedor hidden. The hashed lines in FIG. 59 shows some examples of possibleways to segment data access and provide different levels of granularity.In FIG. 59, monitor mode is shown by block 500 and is the most secure ofall the modes and controls switching between secure and non-secureworlds. Below monitor mode 500 there is a supervisor mode, thiscomprises secure supervisor mode 510 and non-secure supervisor mode 520.Then there is non-secure user mode having applications 522 and 524 andsecure user mode with applications 512, 514 and 516. The monitoringmodes (debug and trace) can be controlled to only monitor non-securemode (to the left of hashed line 501). Alternatively the non-securedomain or world and the secure user mode may be allowed to be monitored(left of 501 and the portion right of 501 that lies below 502). In afurther embodiment the non-secure world and certain applications runningin the secure user domain may be allowed, in this case furthersegmentation by hashed lines 503 occurs. Such divisions help preventleakage of secure data between different users who may be running thedifferent applications. In some controlled cases monitoring of theentire system may be allowed. According to the granularity required thefollowing parts of the core need to have their access controlled duringmonitoring functions.

[0575] There are four registers that can be set on a Debug event; theinstruction Fault Status Register (IFSR), Data Fault Status Register(DFSR), Fault Address Register (FAR), and Instruction Fault AddressRegister (IFAR). These registers should be flushed in some embodimentswhen going from secure world to non-secure world to avoid any leak ofdata.

[0576] PC sample register: The Debug TAP can access the PC through scanchain 7. When debugging in secure world, that value may be maskeddepending on the debug granularity chosen in secure world. It isimportant that non-secure world, or non-secure world plus secure userapplications cannot get any value of the PC while the core is running inthe secure world.

[0577] TLB entries: Using CP15 it is possible to read micro TLB entriesand read and write main TLB entries. We can also control main TLB andmicro TLB loading and matching. This kind of operation must be strictlycontrolled, particularly if secure thread-aware debug requiresassistance of the MMU/MPU.

[0578] Performance Monitor Control register: The performance controlregister gives information on the cache misses, micro TLB misses,external memory requests, branch instruction executed, etc. Non-secureworld should not have access to this data, even in Debug State. Thecounters should be operable in secure world even if debug is disabled insecure world.

[0579] Debugging in cache system: Debugging must be non-intrusive in acached system. It is important is to keep coherency between cache andexternal memory. The Cache can be invalidated using CP15, or the cachecan be forced to be write-through in all regions. In any case, allowingthe modification of cache behaviour in debug can be a security weaknessand should be controlled.

[0580] Endianness: Non-secure world or secure user applications that canaccess to debug should not be allowed to change endianness. Changing theendianness could cause the secure kernel to malfunction. Endiannessaccess is prohibited in debug, according to the granularity.

[0581] Access of the monitoring functions to portions of the core can becontrolled at initiation of the monitoring function. Debug and trace areinitialised in a variety of ways. Embodiments of the present techniquecontrol the access of the monitoring function to certain secure portionsof the core by only allowing initialisation under certain conditions.

[0582] Embodiments of the present technique seek to restrict entry intomonitoring functions with the following granularity:

[0583] By controlling separately intrusive and observable (trace) debug;

[0584] By allowing debug entry in secure user mode only or in the wholesecure world;

[0585] By allowing debug in secure user mode only and moreover takingaccount of the thread ID (application running).

[0586] In order to control the initiation of a monitoring function it isimportant to be aware of how the functions can be initiated. FIG. 60shows a table illustrating the possible ways of initiating a monitoringfunction, the type of monitoring function that is initiated and the waythat such an initiation instruction can be programmed.

[0587] Generally, these monitoring instructions can be entered viasoftware or via hardware, i.e. via the JTAG controller. In order tocontrol the initiation of monitoring functions, control values are used.These comprise enable bits which are condition dependent and thus, if aparticular condition is present, monitoring is only allowed to start ifthe enable bit is set. These bits are stored on a secure register CP14(debug and status control register, DSCR), which is located in ICE 530(see FIG. 67).

[0588] In a preferred embodiment there are four bits that enable/disableintrusive and observable debug, these comprise a secure debug enablebit, a secure trace enable bit, a secure user-mode enable bit and asecure thread aware enable bit. These control values serve to provide adegree of controllable granularity for the monitoring function and assuch can help stop leakage of data from a particular domain. FIG. 61provides a summary of these bits and how they can be accessed.

[0589] These control bits are stored in a register in the secure domainand access to this register is limited to three possibilities. Softwareaccess is provided via ARM coprocessor MRC/MCR instructions and theseare only allowed from the secure supervisor mode. Alternatively,software access can be provided from any other mode with the use of anauthentication code. A further alternative relates more to hardwareaccess and involves the instructions being written via an input port onthe JTAG. In addition to being used to input control values relating tothe availability of monitoring functions, this input port can be used toinput control values relating to other functions of the processor.

[0590] Further details relating to the scan chain and JTAG are givenbelow.

[0591] Register Logic Cell

[0592] Every integrated circuit (IC) consists of two kind of logic:

[0593] Combinatory logic cells; like AND, OR, NV gates. Such gates orcombination of such gates is used to calculate Boolean expressionsaccording to one or several input signals.

[0594] Register logic cells; like LATCH, FLIP-FLOP. Such cells are usedto memorize any signal value. FIG. 62 shows a positive-edge triggeredFLIP-FLOP view:

[0595] When positive-edge event occurs on the clock signal (CK), theoutput (Q) received the value of the input (D); otherwise the output (Q)keeps its value in memory.

[0596] Scan Chain Cell

[0597] For test or debug purpose, it is required to bypass functionalaccess of register logic cells and to have access directly to thecontents of the register logic cells. Thus register cells are integratedin a scan chain cell as shown in FIG. 63.

[0598] In functional mode, SE (Scan Enable) is clear and the registercell works as a single register cell. In test or debug mode, SE is setand input data can come from SI input (Scan In) instead of D input.

[0599] Scan Chain

[0600] All scan chain cells are chained in scan chain as shown in FIG.64.

[0601] In functional mode, SE is clear and all register cells can beaccessed normally and interact with other logic of the circuit. In Testor Debug mode, SE is set and all registers are chained between eachother in a scan chain. Data can come from the first scan chain cell andcan be shifted through any other scan chain cell, at the cadence of eachclock cycle. Data can be shifted out also to see the contents of theregisters.

[0602] TAP Controller

[0603] A debug TAP controller is used to handle several scan chains. TheTAP controller can select a particular scan chain: it connects “Scan In”and “Scan Out” signals to that particular scan-chain. Then data can bescanned into the chain, shifted, or scanned out. The TAP controller iscontrolled externally by a JTAG port interface. FIG. 65 schematicallyillustrates a TAP controller

[0604] JTAG Selective Disable Scan Chain Cell

[0605] For security reasons, some registers might not be accessible byscan chain, even in debug or test mode. A new input called JADI (JTAGAccess Disable) can allow removal dynamically or statically of a scanchain cell from a whole scan chain, without modifying the scan chainstructure in the integrated circuit. FIGS. 66A and B schematically showthis input.

[0606] If JADI is inactive (JADI=0), whether in functional or test ordebug mode, the scan chain works as usual. If JADI is active (JADI=1),and if we are in test or debug mode, some scan chain cells (chosen bydesigner), may be “removed” from the scan chain structure. In order tokeep the same number of scan-chain cell, the JTAG Selective Disable ScanChain Cell use a bypass register. Note that Scan Out (SO) and scan chaincell output (Q) are now different.

[0607]FIG. 67 schematically shows the processor including parts of theJTAG. In normal operation instruction memory 550 communicates with thecore and can under certain circumstances also communicate with registerCP14 and reset the control values. This is generally only allowable fromsecure supervisor mode.

[0608] When debug is initiated instructions are input via debug TAP 580and it is these that control the core. The core in debug runs in a stepby step mode. Debug TAP has access to CP14 via the core (in dependenceupon an access control signal input on the JSDAEN pin shown as JADI pin,JTAG ACCESS DISABLE INPUT in FIG. 45) and the control values can also bereset in this way.

[0609] Access to the CP14 register via debug TAP 580 is controlled by anaccess control signal JSDAEN. This is arranged so that in order foraccess and in particular write access to be allowed JSDAEN must be sethigh. During board stage when the whole processor is being verified,JSDAEN is set high and debug is enabled on the whole system. Once thesystem has been checked, the JSDAEN pin can be tied to ground, thismeans that access to the control values that enable debug in secure modeis now not available via Debug TAP 580. Generally processors inproduction mode have JSDAEN tied to ground. Access to the control valuesis thus, only available via the software route via instruction memory550. Access via this route is limited to secure supervisor mode or toanother mode provided an authentication code is given (see FIG. 68).

[0610] It should be noted that by default debug (intrusive andobservable—trace) are only available in non-secure world. To enable themto be available in secure world the control value enable bits need to beset.

[0611] The advantages of this are that debug can always be initiated byusers to run in non-secure world. Thus, although access to secure worldis not generally available to users in debug this may not be a problemin many cases because access to this world is limited and secure worldhas been fully verified at board stage prior to being made available. Itis therefore foreseen that in many cases debugging of the secure worldwill not be necessary. A secure supervisor can still initiate debug viathe software route of writing CP14 if necessary.

[0612]FIG. 68 schematically shows the control of debug initialisation.In this Figure a portion of the core 600 comprises a storage element 601(which may be a CP15 register as previously discussed) in which isstored a secure status bit S indicative of whether the system is insecure world or not. Core 600 also comprises a register 602 comprisingbits indicative of the mode that the processor is running in, forexample user mode, and a register 603 providing a context identifierthat identifies the application or thread that is currently running onthe core.

[0613] When a breakpoint is reached comparator 610, which compares abreakpoint stored on register 61 lwith the address of the core stored inregister 612, sends a signal to control logic 620. Control logic 620looks at the secure state S, the mode 602 and the thread (contextidentifier) 603 and compares it with the control values and conditionindicators stored on register CP14. If the system is not operating insecure world, then a “enter debug” signal will be output at 630. Ifhowever, the system is operating in secure world, the control logic 620will look at the mode 602, and if it is in user mode will check to seeif user mode enable and debug enable bits are set. If they are thendebug will be initialised provided that a thread aware bit has not beeninitialised. The above illustrates the hierarchical nature of thecontrol values.

[0614] The thread aware portion of the monitoring control is also shownschematically in FIG. 68 along with how the control value stored inregister CP14 can only be changed from secure supervisor mode (in thisembodiment the processor is in production stage and JSDAEN is tied toground). From a secure user mode, secure supervisor mode can be enteredusing an authentication code and then the control value can be set inCP14.

[0615] Control logic 620 outputs an “enter debug” signal when addresscomparator 610 indicates that a breakpoint has been reached providedthread comparator 640 shows that debug is allowable for that thread.This assumes that the thread aware initialisation bit is set in CP14. Ifthe thread aware initialisation bit is set following a breakpoint, debugor trace can only be entered if address and context identifiers matchthose indicated in the breakpoint and in the allowable thread indicator.Following initiation of a monitoring function, the capture of diagnosticdata will only continue while the context identifier is detected bycomparator 640 as an allowed thread. When a context identifier showsthat the application running is not an allowed one, then the capture ofdiagnostic data is suppressed.

[0616] It should be noted that in the preferred embodiment, there issome hierarchy within the granularity. In effect the secure debug ortrace enable bit is at the top, followed by the secure user-mode enablebit and lastly comes the secure thread aware enable bit. This isillustrated in FIGS. 69A and 69B (see below).

[0617] The control values held in the “Debug and Status Control”register (CP14) control secure debug granularity according to thedomain, the mode and the executing thread. It is on top of securesupervisor mode. Once the “Debug and Status Control” register CP14 isconfigured, it's up to secure supervisor mode to program thecorresponding breakpoints, watchpoints, etc to make the core enter DebugState.

[0618]FIG. 69A shows a summary of the secure debug granularity forintrusive debug. Default values at reset are represented in grey colour.

[0619] It is the same for debug granularity concerning observable debug.FIG. 69B shows a summary of secure debug granularity in this case, heredefault values at reset are also represented in grey colour.

[0620] Note that Secure user-mode debug enable bit and Securethread-aware debug enable bit are commonly used for intrusive andobservable debug.

[0621] A thread aware initialisation bit is stored in register CP14 andindicates if granularity by application is required. If the thread awarebit has been initialised, the control logic will further check that theapplication identifier or thread 603 is one indicated in the threadaware control value, if it is, then debug will be initialised. If eitherof the user mode or debug enable bits are not set or the thread awarebit is set and the application running is not one indicated in thethread aware control value, then the breakpoint will be ignored and thecore will continue doing what it was doing and debug will not beinitialised.

[0622] In addition to controlling initialisation of monitoringfunctions, the capture of diagnostic data during a monitor function canalso be controlled in a similar way. In order to do this the core mustcontinue to consider both the control values, i.e. the enable bitsstored in register CP14 and the conditions to which they relate duringoperation of the monitoring function.

[0623]FIG. 70 shows schematically granularity of a monitoring functionwhile it is running. In this case region A relates to a region in whichit is permissible to capture diagnostic data and region B relates toregion in which control values stored in CP14 indicate that it is notpossible to capture diagnostic data.

[0624] Thus, when debug is running and a program is operating in regionA, diagnostic data is output in a step-by-step fashion during debug.When operation switches to Region B, where the capture of diagnosticdata is not allowed, debug no longer proceeds in a step by step fashion,rather it proceeds atomically and no data is captured. This continuesuntil operation of the program re-enters region A whereupon the captureof diagnostic data starts again and debug continues running in astep-by-step fashion.

[0625] In the above embodiment, if secure domain is not enabled, a SMIinstruction is always seen as an atomic event and the capture ofdiagnostic data is suppressed.

[0626] Furthermore, if the thread aware initialisation bit is set thengranularity of the monitoring function during operation with respect toapplication also occurs.

[0627] With regard to observable debug or trace, this is done by ETM andis entirely independent of debug. When trace is enabled ETM works asusual and when it is disabled, ETM hides trace in the secure world, orpart of the secure world depending on the granularity chosen. One way toavoid ETM capturing and tracing diagnostic data in the secure domainwhen this is not enabled is to stall ETM when the S-bit is high. Thiscan be done by combining the S-bit with the ETMPWRDOWN signal, so thatthe ETM values are held at their last values when the core enters secureworld. The ETM should thus trace a SMI instruction and then be stalleduntil the core returns to non-secure world. Thus, the ETM would only seenon-secure activity.

[0628] A summary of some of the different monitoring functions and theirgranularity is given below.

Intrusive Debug at Board Stage

[0629] At board stage when the JSDAEN pin is not tied, there is theability to enable debug everywhere before starting any boot session.Similarly, if we are in secure supervisor mode we have similar rights.

[0630] If we initialise debug in halt debug mode all registers areaccessible (non-secure and secure register banks) and the whole memorycan be dumped, except the bits dedicated to control debug.

[0631] Debug halt mode can be entered from whatever mode and fromwhatever domain. Breakpoints and watchpoints can be set in secure or innon-secure memory. In debug state, it is possible to enter secure worldby simply changing the S bit via an MCR instruction.

[0632] As debug mode can be entered when secure exceptions occur, thevector trap register is extended with new bits which are;

[0633] SMI vector trapping enable

[0634] Secure data abort vector trapping enable

[0635] Secure prefetch abort vector trapping enable

[0636] Secure undefined vector trapping enable.

[0637] In monitor debug mode, if we allow debug everywhere, even when anSMI is called in non-secure world, it is possible to enter secure worldin step-by-step debug. When a breakpoint occurs in secure domain, thesecure abort handler is operable to dump secure register bank and securememory.

[0638] The two abort handlers in secure and in non-secure world givetheir information to the debugger application so that debugger window(on the associated debug controlling PC) can show the register state inboth secure and non-secure worlds.

[0639]FIG. 71A shows what happens when the core is configured in monitordebug mode and debug is enabled in secure world. FIG. 71B shows whathappens when the core is configured in monitor debug mode and the debugis disabled in secure world. This later process will be described below.

Intrusive Debug at Production Stage

[0640] In production stage when JSDAEN is tied and debug is restrictedto non-secure world, unless the secure supervisor determines otherwise,then the table shown in FIG. 71B shows what happens. In this case SMIshould always be considered as an atomic instruction, so that securefunctions are always finished before entering debug state.

[0641] Entering debug halt mode is subject to the followingrestrictions:

[0642] External debug request or internal debug request is taken intoaccount in non-secure world only. If EDBGRQ (external debug request) isasserted while in secure world, the core enters debug halt mode oncesecure function is terminated and the core is returned in non-secureworld.

[0643] Programming a breakpoint or watchpoint on secure memory has noeffect and the core is not stopped when the programmed address matches.

[0644] Vector Trap Register (details of this are given below) concernsnon-secure exceptions only. All extended trapping enable bits explainedbefore have no effect.

[0645] Once in halt debug mode the following restrictions apply:

[0646] S bit cannot be changed to force secure world entry, unlesssecure debug is enabled.

[0647] Mode bits can not be changed if debug is permitted in securesupervisor mode only.

[0648] Dedicated bits that control secure debug cannot be changed.

[0649] If a SMI is loaded and executed (with system speed access), thecore re-enters debug state only when secure function is completelyexecuted.

[0650] In monitor debug mode because monitoring cannot occur in secureworld, the secure abort handler does not need to support a debug monitorprogramme. In non secure world, step-by-step is possible but whenever anSMI is executed secure function is executed entirely in other words anXWSI only “step-over” is allowed while “step-in” and “step-over” arepossible on all other instructions. XWSI is thus considered an atomicinstruction.

[0651] Once secure debug is disabled, we have the followingrestrictions:

[0652] Before entering monitor mode:

[0653] Breakpoints and watchpoints are only taken into account innon-secure world. If bit S is set, breakpoints/watchpoints are bypassed.Note that watchpoints units are also accessible with MCR/MRC (CP14)which is not a security issue as breakpoint/watchpoint has no effect insecure memory.

[0654] BKPT are normally used to replace the instruction on whichbreakpoint is set. This supposes to overwrite this instruction in memoryby BKPT instruction, which will be possible only in non-secure mode.

[0655] Vector Trap Register concerns non-secure exceptions only. Allextended trapping enable bits explained before have no effect. Dataabort and Pre-fetch abort enable bits should be disabled to avoid theprocessor being forced in to an unrecoverable state.

[0656] Via JTAG, we have the same restrictions as for halt mode (S bitcannot be modified, etc)

[0657] Once in monitor mode (non-secure abort mode)

[0658] The non-secure abort handler can dump non-secure world and has novisibility on secure banked registers as well as secure memory.

[0659] Executes secure functions with atomic SMI instruction

[0660] S bit cannot be changed to force secure world entry.

[0661] Mode bits can not be changed if debug is permitted in securesupervisor mode only.

[0662] Note that if an external debug request (EDBGRQ) occurs,

[0663] In non-secure world, the core terminates the current instructionand enters then immediately debug state (in halt mode).

[0664] In secure world, the core terminates the current function andenters the Debug State when it has returned in non-secure world.

[0665] The new debug requirements imply some modifications in corehardware. The S bit must be carefully controlled, and the secure bitmust not be inserted in a scan chain for security reason.

[0666] In summary, in debug, mode bits can be altered only if debug isenabled in secure supervisor mode. It will prevent anybody that hasaccess to debug in the secure domain to have access to all secure worldby altering the system (modifying TBL entries, etc). In that way eachthread can debug its own code, and only its own code. The secure kernelmust be kept safe. Thus when entering debug while the core is running innon-secure world, mode bits can only be altered as before.

[0667] Embodiments of the technique use a new vector trap register. Ifone of the bits in this register is set high and the correspondingvector triggers, the processor enters debug state as if a breakpoint hasbeen set on an instruction fetch from the relevant exception vector. Thebehaviour of these bits may be different according to the value of‘Debug in Secure world Enable’ bit in debug control register.

[0668] The new vector trap register comprises the following bits:D_s_abort, P_s_abort, S_undef, SMI, FIQ, IRQ, Unaligned, D_abort,P_bort, SWI and Undef.

[0669] D_abort bit: should only be set when debug is enabled in secureworld and when debug is configured in halt debug mode. In monitor debugmode, this bit should never bit set. If debug in secure world isdisabled, this bit has no effect whatever its value.

[0670] P_abort bit: same as D_abort bit.

[0671] S_undef bit: should only be set when debug is enable in secureworld. If debug in secure world is disabled, this bit has no effectwhatever its value is.

[0672] SMI bit: should only be set when debug is enabled in secureworld. If debug in secure world is disabled, this bit has no effectwhatever its value is.

[0673] FIQ, IRQ, D_abort, P_abort, SWI, undef bits: correspond tonon-secure exceptions, so they are valid even if debug in secure worldis disabled. Note that D_abort and P_abort should not be asserted highin monitor mode.

[0674] Reset bit: as we enter secure world when reset occurs, this bitis valid only when debug in secure world is enabled, otherwise it has noeffect.

[0675] Although a particular embodiment of the invention has beendescribed herein, it will be apparent that the invention is not limitedthereto, and that many modifications and additions may be made withinthe scope of the invention. For example, various combinations of thefeatures of the following dependent could be made with the features ofthe independent claims without departing from the scope of the presentinvention.

We claim:
 1. A data processing apparatus, comprising: a processoroperable in a plurality of modes and a plurality of domains, saidplurality of domains comprising a secure domain and a non-secure domain,said plurality of modes including at least one non-secure mode being amode in the non-secure domain, and at least one secure mode being a modein the secure domain, said processor being operable such that whenexecuting a program in a secure mode said program has access to securedata which is not accessible when said processor is operating in anon-secure mode; a memory operable to store data required by theprocessor and comprising secure memory for storing secure data andnon-secure memory for storing non-secure data, the processor beingoperable to issue a memory access request when access to an item of datain the memory is required; at least one memory management unit operable,upon receipt of the memory access request from the processor, to performconversion of a virtual address specified by the memory access requestto a physical address; a first set of tables, each table in the firstset containing a number of first descriptors, each first descriptorcontaining at least a virtual address portion and a correspondingintermediate address portion; a second set of tables, each table in thesecond set containing a number of second descriptors, each seconddescriptor containing at least an intermediate address portion and acorresponding physical address portion, the second set of tables beingmanaged by the processor when operating in a privileged mode which isnot a non-secure mode; the at least one memory management unit beingoperable to cause predetermined tables in said first and second set tobe referenced to enable the conversion of the virtual address specifiedby the memory access request to a physical address.
 2. A data processingapparatus as claimed in claim 1, wherein the privileged mode is amonitor mode in which the processor is operable to manage switchingbetween said secure domain and said non-secure domain.
 3. A dataprocessing apparatus as claimed in claim 1, wherein said privileged modeis a privileged secure mode.
 4. A data processing apparatus as claimedin claim 3, wherein in said at least one non-secure mode the processoris operable under the control of a non-secure operating system and insaid at least one secure mode the processor is operable under thecontrol of a secure operating system, the secure operating system beingoperable to manage the second set of tables when the processor isoperating in the privileged secure mode.
 5. A data processing apparatusas claimed in claim 2, wherein when switching between the secure domainand the non-secure domain, the processor is operable in the monitor modeto select the predetermined tables in said first and second sets,dependent on whether the domain being switched to is the secure domainor the non-secure domain.
 6. A data processing apparatus as claimed inclaim 1, wherein the predetermined tables within said first and secondsets are selected when the tables are to be referenced dependent onwhether the processor is operating in a secure mode or a non-secure modeat the time the memory access request is issued.
 7. A data processingapparatus as claimed in claim 1, wherein the at least one memorymanagement unit comprises a first memory management unit and a secondmemory management unit.
 8. A data processing apparatus as claimed inclaim 7, wherein the tables in the first set are associated with thefirst memory management unit and the tables in the second set areassociated with the second memory management unit.
 9. A data processingapparatus as claimed in claim 7, wherein if the first memory managementunit needs to access a first descriptor within a predetermined table ofsaid first set, it issues a table lookup request specifying anintermediate address for that first descriptor, the second memorymanagement unit being operable to receive the table lookup request anddetermine the physical address corresponding to that intermediateaddress.
 10. A data processing apparatus as claimed in claim 9, whereinthe second memory management unit is then operable to cause the firstdescriptor at that physical address to be retrieved and returned to thefirst memory management unit.
 11. A data processing apparatus as claimedin claim 7, wherein the first memory management unit comprises a firstinternal storage unit for storing first descriptors retrieved from thepredetermined table of the first set, and used by the first memorymanagement unit to derive access control information used to perform theconversion of the virtual address into a corresponding intermediateaddress.
 12. A data processing apparatus as claimed in claim 11, whereinthe first internal storage unit is a first translation lookaside buffer(TLB) operable to store the first descriptors retrieved from thepredetermined table of the first set.
 13. A data processing apparatus asclaimed in claim 12, wherein the first TLB is a first main TLB forstoring the first descriptors retrieved by the first memory managementunit from the predetermined table of the first set, and the internalstorage further comprises a micro-TLB for storing the access controlinformation derived from the first descriptors, the access controlinformation comprising conversions between a number of virtual addressportions and corresponding intermediate address portions, and the accesscontrol information being transferred from the first main TLB to themicro-TLB prior to use of that access control information by the firstmemory management unit.
 14. A data processing apparatus as claimed inclaim 7, wherein the first memory management unit comprises a firstinternal storage unit for storing new descriptors derived fromcorresponding first and second descriptors retrieved from thepredetermined tables of the first and second sets, and used by the firstmemory management unit to derive access control information used toperform the conversion of the virtual address into a correspondingphysical address.
 15. A data processing apparatus as claimed in claim14, wherein the first internal storage unit is a first translationlookaside buffer (TLB) operable to store the new descriptors derivedfrom corresponding first and second descriptors.
 16. A data processingapparatus as claimed in claim 15, wherein the first TLB is a first mainTLB for storing the new descriptors derived from corresponding first andsecond descriptors, and the internal storage unit further comprises amicro-TLB for storing the access control information, the access controlinformation comprising conversions between a number of virtual addressportions and corresponding physical address portions, and the accesscontrol information being transferred from the first main TLB to themicro-TLB prior to use of that access control information by the firstmemory management unit.
 17. A data processing apparatus as claimed inclaim 7, wherein the second memory management unit comprises a secondinternal storage unit for storing second descriptors retrieved from thepredetermined table of the second set, and used by the second memorymanagement unit to derive access control information used to perform theconversion of the intermediate address into a corresponding physicaladdress.
 18. A data processing apparatus as claimed in claim 17, whereinthe second internal storage unit is a second translation lookasidebuffer (TLB) operable to store the second descriptors retrieved from thepredetermined table of the second set.
 19. A data processing apparatusas claimed in claim 18, wherein the second TLB is a second main TLB forstoring the second descriptors retrieved by the second memory managementunit from the predetermined table of the second set, and the internalstorage further comprises a micro-TLB for storing the access controlinformation derived from the second descriptors, the access controlinformation comprising conversions between a number of intermediateaddress portions and corresponding physical address portions, and theaccess control information being transferred from the second main TLB tothe micro-TLB prior to use of that access control information by thesecond memory management unit.
 20. A data processing apparatus asclaimed in claim 18, wherein the first internal storage unit is a firsttranslation lookaside buffer (TLB) operable to store the firstdescriptors retrieved from the predetermined table of the first set, andwherein the first and second sets of tables each comprise at least asecure table and a non-secure table, the first and second TLBscomprising a flag associated with each descriptor stored therein toidentify whether that descriptor is derived from said non-secure tableor said secure table.
 21. A data processing apparatus as claimed inclaim 19, wherein the first and second sets of tables each comprise atleast a secure table and a non-secure table, the first and second TLBscomprising a flag associated with each descriptor stored therein toidentify whether that descriptor is derived from said non-secure tableor said secure table, and wherein the micro-TLB of both the first andsecond memory management units is flushed whenever the mode of operationof the processor changes between a secure mode and a non-secure mode, inthe secure mode access control information only being transferred to themicro-TLB from a descriptor in the associated first or second main TLBthat said associated flag indicates is from the secure table, and in thenon-secure mode access control information only being transferred to themicro-TLB from a descriptor in the associated first or second main TLBthat said associated flag indicates is from the non-secure table.
 22. Adata processing apparatus as claimed claim 1, wherein the at least onememory management unit comprises a single memory management unit, andthe processor is operable to execute table merging code to reference thepredetermined tables of the first and second sets in order to producefrom a first descriptor and an associated second descriptor a newdescriptor associating a virtual address portion with a correspondingphysical address portion.
 23. A data processing apparatus as claimed inclaim 22, wherein the table merging code is operable to retrieve thefirst descriptor after referencing the predetermined table in the secondset to obtain the physical address of the first descriptor.
 24. A dataprocessing apparatus as claimed in claim 23, wherein the table mergingcode is operable to use the first descriptor to determine theintermediate address corresponding to the virtual address specified bythe memory access request, and to then reference the predetermined tablein the second set to obtain the second descriptor providing a physicaladdress for that intermediate address, whereafter the table merging codeis operable to merge the first and second descriptors to produce the newdescriptor.
 25. A data processing apparatus as claimed in claim 22,wherein the single memory management unit comprises an internal storageunit for storing the new descriptor produced by the table merging code,and used by the single memory management unit to derive access controlinformation used to perform the conversion of the virtual address into acorresponding physical address.
 26. A data processing apparatus asclaimed in claim 25, wherein the processor is operable to execute thetable merging code when the access control information required todetermine the physical address for the memory access request is notfound in the internal storage unit.
 27. A data processing apparatus asclaimed in claim 25, wherein the internal storage unit is a translationlookaside buffer (TLB) operable to store the new descriptors produced bythe table merging code.
 28. A data processing apparatus as claimed inclaim 27, wherein the TLB is a main TLB for storing the new descriptorsobtained by the single memory management unit from the table mergingcode, and the internal storage further comprises a micro-TLB for storingthe access control information derived from the new descriptors, theaccess control information comprising conversions between a number ofvirtual address portions and corresponding physical address portions,and the access control information being transferred from the main TLBto the micro-TLB prior to use of that access control information by thesingle memory management unit.
 29. A data processing apparatus asclaimed in claim 27, wherein the first and second sets of tables eachcomprise at least a secure table and a non-secure table, the TLBcomprising a flag associated with each new descriptor stored therein toidentify whether that new descriptor is derived from said non-securetables or said secure tables.
 30. A data processing apparatus as claimedin claim 28, wherein the first and second sets of tables each compriseat least a secure table and a non-secure table, the TLB comprising aflag associated with each new descriptor stored therein to identifywhether that new descriptor is derived from said non-secure tables orsaid secure tables, and wherein the micro-TLB of the single memorymanagement unit is flushed whenever the mode of operation of theprocessor changes between a secure mode and a non-secure mode, in thesecure mode access control information only being transferred to themicro-TLB from a new descriptor in the main TLB that said associatedflag indicates is derived from secure tables, and in the non-secure modeaccess control information only being transferred to the micro-TLB froma new descriptor in the main TLB that said associated flag indicates isderived from non-secure tables.
 31. A data processing apparatus asclaimed in claim 22, wherein the privileged mode is a monitor mode inwhich the processor is operable to manage switching between said securedomain and said non-secure domain, and wherein the table merging code isexecuted by the processor when operating in the monitor mode.
 32. A dataprocessing apparatus as claimed in claim 1, wherein said first andsecond sets of tables comprise page tables.
 33. A data processingapparatus as claimed in claim 1, wherein the first set of tables and thesecond set of tables are stored within said memory.
 34. A method ofcontrolling access to a memory in a data processing apparatus, the dataprocessing apparatus comprising a processor operable in a plurality ofmodes and a plurality of domains, said plurality of domains comprising asecure domain and a non-secure domain, said plurality of modes includingat least one non-secure mode being a mode in the non-secure domain, andat least one secure mode being a mode in the secure domain, saidprocessor being operable such that when executing a program in a securemode said program has access to secure data which is not accessible whensaid processor is operating in a non-secure mode, the memory beingoperable to store data required by the processor and comprising securememory for storing secure data and non-secure memory for storingnon-secure data, the method comprising the steps of: providing a firstset of tables, each table in the first set containing a number of firstdescriptors, each first descriptor containing at least a virtual addressportion and a corresponding intermediate address portion; providing asecond set of tables, each table in the second set containing a numberof second descriptors, each second descriptor containing at least anintermediate address portion and a corresponding physical addressportion, the second set of tables being managed by the processor whenoperating in a privileged mode which is not a non-secure mode; issuingfrom the processor a memory access request when access to an item ofdata in the memory is required; and performing conversion of a virtualaddress specified by the memory access request to a physical addresswith reference to predetermined tables in said first and second set. 35.A method as claimed in claim 34, wherein the privileged mode is amonitor mode in which the processor is operable to manage switchingbetween said secure domain and said non-secure domain.
 36. A method asclaimed in claim 34, wherein said privileged mode is a privileged securemode.
 37. A method as claimed in claim 36, wherein in said at least onenon-secure mode the processor is operable under the control of anon-secure operating system and in said at least one secure mode theprocessor is operable under the control of a secure operating system,the secure operating system being operable to manage the second set oftables when the processor is operating in the privileged secure mode.38. A method as claimed in claim 35, wherein when switching between thesecure domain and the non-secure domain, the processor is operable inthe monitor mode to select the predetermined tables in said first andsecond sets, dependent on whether the domain being switched to is thesecure domain or the non-secure domain.
 39. A method as claimed in claim34, wherein the predetermined tables within said first and second setsare selected when the tables are to be referenced dependent on whetherthe processor is operating in a secure mode or a non-secure mode at thetime the memory access request is issued.
 40. A method as claimed inclaim 34, wherein said step of performing conversion of a virtualaddress to a physical address is performed by at least one of a firstmemory management unit and a second memory management unit.
 41. A methodas claimed in claim 40, wherein if the first memory management unitneeds to access a first descriptor within a predetermined table of saidfirst set, the method further comprises the steps of: issuing from thefirst memory management unit a table lookup request specifying anintermediate address for that first descriptor; and receiving the tablelookup request at the second memory management unit and determining thephysical address corresponding to that intermediate address.
 42. Amethod as claimed in claim 41, further comprising the step of: causingthe first descriptor at that physical address to be retrieved andreturned to the first memory management unit.
 43. A method as claimed inclaim 42, further comprising the steps of: causing a second descriptorwithin a predetermined table of said second set to be retrieved; andmerging the first descriptor and second descriptor in order to produce anew descriptor for storing in the first memory management unit, the newdescriptor containing at least a virtual address portion and acorresponding physical address portion.
 44. A method as claimed in claim34, wherein the data processing apparatus comprises a single memorymanagement unit, and the method comprises the step of: executing tablemerging code to reference the predetermined tables of the first andsecond sets in order to produce from a first descriptor and anassociated second descriptor a new descriptor associating a virtualaddress portion with a corresponding physical address portion.
 45. Amethod as claimed in claim 44, wherein the table merging code isoperable to perform the steps of: referencing the predetermined table inthe second set to obtain the physical address of the first descriptor;and retrieving the first descriptor.
 46. A method as claimed in claim45, wherein the table merging code is further operable to perform thesteps of: using the first descriptor to determine the intermediateaddress corresponding to the virtual address specified by the memoryaccess request; referencing the predetermined table in the second set toobtain the second descriptor providing a physical address for thatintermediate address; and merging the first and second descriptors toproduce the new descriptor.
 47. A method as claimed in claim 44, whereinthe single memory management unit comprises an internal storage unit forstoring access control information derived from the new descriptorproduced by the table merging code, and used by the single memorymanagement unit to perform the conversion of the virtual address into acorresponding physical address.
 48. A method as claimed in claim 47,wherein the processor is operable to execute the table merging code whenthe access control information required to determine the physicaladdress for the memory access request is not found in the internalstorage unit.
 49. A method as claimed in claim 44, wherein theprivileged mode is a monitor mode in which the processor is operable tomanage switching between said secure domain and said non-secure domain,and wherein the table merging code is executed by the processor whenoperating in the monitor mode.
 50. A computer program providing tablemerging code and operable to configure a processor of a data processingapparatus to perform the method of claim
 44. 51. A computer programproduct carrying a computer program as claimed in claim 50.